ata: libata-scsi: fix ata_dsm_trim_pages() kernel-docgrafted

make htmldocs fails with:

  Documentation/driver-api/libata:604: ./drivers/ata/libata-scsi.c:2301:
  ERROR: Unexpected indentation. [docutils]

The Return: section of ata_dsm_trim_pages() ends its first sentence with
a colon and continues with an indented bullet list. reStructuredText
requires a blank line before an indented block, so docutils chokes on
the list.

Simply adding the missing blank line does not work either: Return: is a
kernel-doc "special section", which is terminated by the first blank
line, so the bullet list would end up in the Description section,
detached from the sentence introducing it.

Spell the two bounds out as prose instead, so that the Return: section
stays self-contained. Documentation-only change.

Fixes: e64e6b5dc867 ("ata: libata-scsi: scale DSM TRIM payload by MAX PAGES PER DSM COMMAND")
Reported-by: Thomas Huth <thuth@redhat.com>
Closes: https://lore.kernel.org/linux-ide/b0a0b8e8-cc4f-4c7b-8bc5-fee0712d405a@redhat.com/
Reviewed-by: Damien Le Moal <dlemoal@kernel.org>
Reviewed-by: Thomas Huth <thuth@redhat.com>
Link: https://lore.kernel.org/r/20260916130031.29990-2-cassel@kernel.org
Signed-off-by: Niklas Cassel <cassel@kernel.org>

5 files changed, 6927 insertions, 0 deletions

diff --git a/Documentation/filesystems/xfs/index.rst b/Documentation/filesystems/xfs/index.rst
new file mode 100644
index 000000000..ab66c57a5
--- /dev/null
+++ b/Documentation/filesystems/xfs/index.rst
@@ -0,0 +1,14 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============================
+XFS Filesystem Documentation
+============================
+
+.. toctree::
+   :maxdepth: 2
+   :numbered:
+
+   xfs-delayed-logging-design
+   xfs-maintainer-entry-profile
+   xfs-self-describing-metadata
+   xfs-online-fsck-design
diff --git a/Documentation/filesystems/xfs/xfs-delayed-logging-design.rst b/Documentation/filesystems/xfs/xfs-delayed-logging-design.rst
new file mode 100644
index 000000000..2a2705e97
--- /dev/null
+++ b/Documentation/filesystems/xfs/xfs-delayed-logging-design.rst
@@ -0,0 +1,1087 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==================
+XFS Logging Design
+==================
+
+Preamble
+========
+
+This document describes the design and algorithms that the XFS journalling
+subsystem is based on. This document describes the design and algorithms that
+the XFS journalling subsystem is based on so that readers may familiarize
+themselves with the general concepts of how transaction processing in XFS works.
+
+We begin with an overview of transactions in XFS, followed by describing how
+transaction reservations are structured and accounted, and then move into how we
+guarantee forwards progress for long running transactions with finite initial
+reservations bounds. At this point we need to explain how relogging works. With
+the basic concepts covered, the design of the delayed logging mechanism is
+documented.
+
+
+Introduction
+============
+
+XFS uses Write Ahead Logging for ensuring changes to the filesystem metadata
+are atomic and recoverable. For reasons of space and time efficiency, the
+logging mechanisms are varied and complex, combining intents, logical and
+physical logging mechanisms to provide the necessary recovery guarantees the
+filesystem requires.
+
+Some objects, such as inodes and dquots, are logged in logical format where the
+details logged are made up of the changes to in-core structures rather than
+on-disk structures. Other objects - typically buffers - have their physical
+changes logged. Long running atomic modifications have individual changes
+chained together by intents, ensuring that journal recovery can restart and
+finish an operation that was only partially done when the system stopped
+functioning.
+
+The reason for these differences is to keep the amount of log space and CPU time
+required to process objects being modified as small as possible and hence the
+logging overhead as low as possible. Some items are very frequently modified,
+and some parts of objects are more frequently modified than others, so keeping
+the overhead of metadata logging low is of prime importance.
+
+The method used to log an item or chain modifications together isn't
+particularly important in the scope of this document. It suffices to know that
+the method used for logging a particular object or chaining modifications
+together are different and are dependent on the object and/or modification being
+performed. The logging subsystem only cares that certain specific rules are
+followed to guarantee forwards progress and prevent deadlocks.
+
+
+Transactions in XFS
+===================
+
+XFS has two types of high level transactions, defined by the type of log space
+reservation they take. These are known as "one shot" and "permanent"
+transactions. Permanent transaction reservations can take reservations that span
+commit boundaries, whilst "one shot" transactions are for a single atomic
+modification.
+
+The type and size of reservation must be matched to the modification taking
+place.  This means that permanent transactions can be used for one-shot
+modifications, but one-shot reservations cannot be used for permanent
+transactions.
+
+In the code, a one-shot transaction pattern looks somewhat like this::
+
+	tp = xfs_trans_alloc(<reservation>)
+	<lock items>
+	<join item to transaction>
+	<do modification>
+	xfs_trans_commit(tp);
+
+As items are modified in the transaction, the dirty regions in those items are
+tracked via the transaction handle.  Once the transaction is committed, all
+resources joined to it are released, along with the remaining unused reservation
+space that was taken at the transaction allocation time.
+
+In contrast, a permanent transaction is made up of multiple linked individual
+transactions, and the pattern looks like this::
+
+	tp = xfs_trans_alloc(<reservation>)
+	xfs_ilock(ip, XFS_ILOCK_EXCL)
+
+	loop {
+		xfs_trans_ijoin(tp, 0);
+		<do modification>
+		xfs_trans_log_inode(tp, ip);
+		xfs_trans_roll(&tp);
+	}
+
+	xfs_trans_commit(tp);
+	xfs_iunlock(ip, XFS_ILOCK_EXCL);
+
+While this might look similar to a one-shot transaction, there is an important
+difference: xfs_trans_roll() performs a specific operation that links two
+transactions together::
+
+	ntp = xfs_trans_dup(tp);
+	xfs_trans_commit(tp);
+	xfs_trans_reserve(ntp);
+
+This results in a series of "rolling transactions" where the inode is locked
+across the entire chain of transactions.  Hence while this series of rolling
+transactions is running, nothing else can read from or write to the inode and
+this provides a mechanism for complex changes to appear atomic from an external
+observer's point of view.
+
+It is important to note that a series of rolling transactions in a permanent
+transaction does not form an atomic change in the journal. While each
+individual modification is atomic, the chain is *not atomic*. If we crash half
+way through, then recovery will only replay up to the last transactional
+modification the loop made that was committed to the journal.
+
+This affects long running permanent transactions in that it is not possible to
+predict how much of a long running operation will actually be recovered because
+there is no guarantee of how much of the operation reached stale storage. Hence
+if a long running operation requires multiple transactions to fully complete,
+the high level operation must use intents and deferred operations to guarantee
+recovery can complete the operation once the first transactions is persisted in
+the on-disk journal.
+
+
+Transactions are Asynchronous
+=============================
+
+In XFS, all high level transactions are asynchronous by default. This means that
+xfs_trans_commit() does not guarantee that the modification has been committed
+to stable storage when it returns. Hence when a system crashes, not all the
+completed transactions will be replayed during recovery.
+
+However, the logging subsystem does provide global ordering guarantees, such
+that if a specific change is seen after recovery, all metadata modifications
+that were committed prior to that change will also be seen.
+
+For single shot operations that need to reach stable storage immediately, or
+ensuring that a long running permanent transaction is fully committed once it is
+complete, we can explicitly tag a transaction as synchronous. This will trigger
+a "log force" to flush the outstanding committed transactions to stable storage
+in the journal and wait for that to complete.
+
+Synchronous transactions are rarely used, however, because they limit logging
+throughput to the IO latency limitations of the underlying storage. Instead, we
+tend to use log forces to ensure modifications are on stable storage only when
+a user operation requires a synchronisation point to occur (e.g. fsync).
+
+
+Transaction Reservations
+========================
+
+It has been mentioned a number of times now that the logging subsystem needs to
+provide a forwards progress guarantee so that no modification ever stalls
+because it can't be written to the journal due to a lack of space in the
+journal. This is achieved by the transaction reservations that are made when
+a transaction is first allocated. For permanent transactions, these reservations
+are maintained as part of the transaction rolling mechanism.
+
+A transaction reservation provides a guarantee that there is physical log space
+available to write the modification into the journal before we start making
+modifications to objects and items. As such, the reservation needs to be large
+enough to take into account the amount of metadata that the change might need to
+log in the worst case. This means that if we are modifying a btree in the
+transaction, we have to reserve enough space to record a full leaf-to-root split
+of the btree. As such, the reservations are quite complex because we have to
+take into account all the hidden changes that might occur.
+
+For example, a user data extent allocation involves allocating an extent from
+free space, which modifies the free space trees. That's two btrees.  Inserting
+the extent into the inode's extent map might require a split of the extent map
+btree, which requires another allocation that can modify the free space trees
+again.  Then we might have to update reverse mappings, which modifies yet
+another btree which might require more space. And so on.  Hence the amount of
+metadata that a "simple" operation can modify can be quite large.
+
+This "worst case" calculation provides us with the static "unit reservation"
+for the transaction that is calculated at mount time. We must guarantee that the
+log has this much space available before the transaction is allowed to proceed
+so that when we come to write the dirty metadata into the log we don't run out
+of log space half way through the write.
+
+For one-shot transactions, a single unit space reservation is all that is
+required for the transaction to proceed. For permanent transactions, however, we
+also have a "log count" that affects the size of the reservation that is to be
+made.
+
+While a permanent transaction can get by with a single unit of space
+reservation, it is somewhat inefficient to do this as it requires the
+transaction rolling mechanism to re-reserve space on every transaction roll. We
+know from the implementation of the permanent transactions how many transaction
+rolls are likely for the common modifications that need to be made.
+
+For example, an inode allocation is typically two transactions - one to
+physically allocate a free inode chunk on disk, and another to allocate an inode
+from an inode chunk that has free inodes in it.  Hence for an inode allocation
+transaction, we might set the reservation log count to a value of 2 to indicate
+that the common/fast path transaction will commit two linked transactions in a
+chain. Each time a permanent transaction rolls, it consumes an entire unit
+reservation.
+
+Hence when the permanent transaction is first allocated, the log space
+reservation is increased from a single unit reservation to multiple unit
+reservations. That multiple is defined by the reservation log count, and this
+means we can roll the transaction multiple times before we have to re-reserve
+log space when we roll the transaction. This ensures that the common
+modifications we make only need to reserve log space once.
+
+If the log count for a permanent transaction reaches zero, then it needs to
+re-reserve physical space in the log. This is somewhat complex, and requires
+an understanding of how the log accounts for space that has been reserved.
+
+
+Log Space Accounting
+====================
+
+The position in the log is typically referred to as a Log Sequence Number (LSN).
+The log is circular, so the positions in the log are defined by the combination
+of a cycle number - the number of times the log has been overwritten - and the
+offset into the log.  A LSN carries the cycle in the upper 32 bits and the
+offset in the lower 32 bits. The offset is in units of "basic blocks" (512
+bytes). Hence we can do relatively simple LSN based math to keep track of
+available space in the log.
+
+Log space accounting is done via a pair of constructs called "grant heads".  The
+position of the grant heads is an absolute value, so the amount of space
+available in the log is defined by the distance between the position of the
+grant head and the current log tail. That is, how much space can be
+reserved/consumed before the grant heads would fully wrap the log and overtake
+the tail position.
+
+The first grant head is the "reserve" head. This tracks the byte count of the
+reservations currently held by active transactions. It is a purely in-memory
+accounting of the space reservation and, as such, actually tracks byte offsets
+into the log rather than basic blocks. Hence it technically isn't using LSNs to
+represent the log position, but it is still treated like a split {cycle,offset}
+tuple for the purposes of tracking reservation space.
+
+The reserve grant head is used to accurately account for exact transaction
+reservations amounts and the exact byte count that modifications actually make
+and need to write into the log. The reserve head is used to prevent new
+transactions from taking new reservations when the head reaches the current
+tail. It will block new reservations in a FIFO queue and as the log tail moves
+forward it will wake them in order once sufficient space is available. This FIFO
+mechanism ensures no transaction is starved of resources when log space
+shortages occur.
+
+The other grant head is the "write" head. Unlike the reserve head, this grant
+head contains an LSN and it tracks the physical space usage in the log. While
+this might sound like it is accounting the same state as the reserve grant head
+- and it mostly does track exactly the same location as the reserve grant head -
+there are critical differences in behaviour between them that provides the
+forwards progress guarantees that rolling permanent transactions require.
+
+These differences when a permanent transaction is rolled and the internal "log
+count" reaches zero and the initial set of unit reservations have been
+exhausted. At this point, we still require a log space reservation to continue
+the next transaction in the sequeunce, but we have none remaining. We cannot
+sleep during the transaction commit process waiting for new log space to become
+available, as we may end up on the end of the FIFO queue and the items we have
+locked while we sleep could end up pinning the tail of the log before there is
+enough free space in the log to fulfill all of the pending reservations and
+then wake up transaction commit in progress.
+
+To take a new reservation without sleeping requires us to be able to take a
+reservation even if there is no reservation space currently available. That is,
+we need to be able to *overcommit* the log reservation space. As has already
+been detailed, we cannot overcommit physical log space. However, the reserve
+grant head does not track physical space - it only accounts for the amount of
+reservations we currently have outstanding. Hence if the reserve head passes
+over the tail of the log all it means is that new reservations will be throttled
+immediately and remain throttled until the log tail is moved forward far enough
+to remove the overcommit and start taking new reservations. In other words, we
+can overcommit the reserve head without violating the physical log head and tail
+rules.
+
+As a result, permanent transactions only "regrant" reservation space during
+xfs_trans_commit() calls, while the physical log space reservation - tracked by
+the write head - is then reserved separately by a call to xfs_log_reserve()
+after the commit completes. Once the commit completes, we can sleep waiting for
+physical log space to be reserved from the write grant head, but only if one
+critical rule has been observed::
+
+	Code using permanent reservations must always log the items they hold
+	locked across each transaction they roll in the chain.
+
+"Re-logging" the locked items on every transaction roll ensures that the items
+attached to the transaction chain being rolled are always relocated to the
+physical head of the log and so do not pin the tail of the log. If a locked item
+pins the tail of the log when we sleep on the write reservation, then we will
+deadlock the log as we cannot take the locks needed to write back that item and
+move the tail of the log forwards to free up write grant space. Re-logging the
+locked items avoids this deadlock and guarantees that the log reservation we are
+making cannot self-deadlock.
+
+If all rolling transactions obey this rule, then they can all make forwards
+progress independently because nothing will block the progress of the log
+tail moving forwards and hence ensuring that write grant space is always
+(eventually) made available to permanent transactions no matter how many times
+they roll.
+
+
+Re-logging Explained
+====================
+
+XFS allows multiple separate modifications to a single object to be carried in
+the log at any given time.  This allows the log to avoid needing to flush each
+change to disk before recording a new change to the object. XFS does this via a
+method called "re-logging". Conceptually, this is quite simple - all it requires
+is that any new change to the object is recorded with a *new copy* of all the
+existing changes in the new transaction that is written to the log.
+
+That is, if we have a sequence of changes A through to F, and the object was
+written to disk after change D, we would see in the log the following series
+of transactions, their contents and the log sequence number (LSN) of the
+transaction::
+
+	Transaction		Contents	LSN
+	   A			   A		   X
+	   B			  A+B		  X+n
+	   C			 A+B+C		 X+n+m
+	   D			A+B+C+D		X+n+m+o
+	    <object written to disk>
+	   E			   E		   Y (> X+n+m+o)
+	   F			  E+F		  Y+p
+
+In other words, each time an object is relogged, the new transaction contains
+the aggregation of all the previous changes currently held only in the log.
+
+This relogging technique allows objects to be moved forward in the log so that
+an object being relogged does not prevent the tail of the log from ever moving
+forward.  This can be seen in the table above by the changing (increasing) LSN
+of each subsequent transaction, and it's the technique that allows us to
+implement long-running, multiple-commit permanent transactions. 
+
+A typical example of a rolling transaction is the removal of extents from an
+inode which can only be done at a rate of two extents per transaction because
+of reservation size limitations. Hence a rolling extent removal transaction
+keeps relogging the inode and btree buffers as they get modified in each
+removal operation. This keeps them moving forward in the log as the operation
+progresses, ensuring that current operation never gets blocked by itself if the
+log wraps around.
+
+Hence it can be seen that the relogging operation is fundamental to the correct
+working of the XFS journalling subsystem. From the above description, most
+people should be able to see why the XFS metadata operations writes so much to
+the log - repeated operations to the same objects write the same changes to
+the log over and over again. Worse is the fact that objects tend to get
+dirtier as they get relogged, so each subsequent transaction is writing more
+metadata into the log.
+
+It should now also be obvious how relogging and asynchronous transactions go
+hand in hand. That is, transactions don't get written to the physical journal
+until either a log buffer is filled (a log buffer can hold multiple
+transactions) or a synchronous operation forces the log buffers holding the
+transactions to disk. This means that XFS is doing aggregation of transactions
+in memory - batching them, if you like - to minimise the impact of the log IO on
+transaction throughput.
+
+The limitation on asynchronous transaction throughput is the number and size of
+log buffers made available by the log manager. By default there are 8 log
+buffers available and the size of each is 32kB - the size can be increased up
+to 256kB by use of a mount option.
+
+Effectively, this gives us the maximum bound of outstanding metadata changes
+that can be made to the filesystem at any point in time - if all the log
+buffers are full and under IO, then no more transactions can be committed until
+the current batch completes. It is now common for a single current CPU core to
+be to able to issue enough transactions to keep the log buffers full and under
+IO permanently. Hence the XFS journalling subsystem can be considered to be IO
+bound.
+
+Delayed Logging: Concepts
+=========================
+
+The key thing to note about the asynchronous logging combined with the
+relogging technique XFS uses is that we can be relogging changed objects
+multiple times before they are committed to disk in the log buffers. If we
+return to the previous relogging example, it is entirely possible that
+transactions A through D are committed to disk in the same log buffer.
+
+That is, a single log buffer may contain multiple copies of the same object,
+but only one of those copies needs to be there - the last one "D", as it
+contains all the changes from the previous changes. In other words, we have one
+necessary copy in the log buffer, and three stale copies that are simply
+wasting space. When we are doing repeated operations on the same set of
+objects, these "stale objects" can be over 90% of the space used in the log
+buffers. It is clear that reducing the number of stale objects written to the
+log would greatly reduce the amount of metadata we write to the log, and this
+is the fundamental goal of delayed logging.
+
+From a conceptual point of view, XFS is already doing relogging in memory (where
+memory == log buffer), only it is doing it extremely inefficiently. It is using
+logical to physical formatting to do the relogging because there is no
+infrastructure to keep track of logical changes in memory prior to physically
+formatting the changes in a transaction to the log buffer. Hence we cannot avoid
+accumulating stale objects in the log buffers.
+
+Delayed logging is the name we've given to keeping and tracking transactional
+changes to objects in memory outside the log buffer infrastructure. Because of
+the relogging concept fundamental to the XFS journalling subsystem, this is
+actually relatively easy to do - all the changes to logged items are already
+tracked in the current infrastructure. The big problem is how to accumulate
+them and get them to the log in a consistent, recoverable manner.
+Describing the problems and how they have been solved is the focus of this
+document.
+
+One of the key changes that delayed logging makes to the operation of the
+journalling subsystem is that it disassociates the amount of outstanding
+metadata changes from the size and number of log buffers available. In other
+words, instead of there only being a maximum of 2MB of transaction changes not
+written to the log at any point in time, there may be a much greater amount
+being accumulated in memory. Hence the potential for loss of metadata on a
+crash is much greater than for the existing logging mechanism.
+
+It should be noted that this does not change the guarantee that log recovery
+will result in a consistent filesystem. What it does mean is that as far as the
+recovered filesystem is concerned, there may be many thousands of transactions
+that simply did not occur as a result of the crash. This makes it even more
+important that applications that care about their data use fsync() where they
+need to ensure application level data integrity is maintained.
+
+It should be noted that delayed logging is not an innovative new concept that
+warrants rigorous proofs to determine whether it is correct or not. The method
+of accumulating changes in memory for some period before writing them to the
+log is used effectively in many filesystems including ext3 and ext4. Hence
+no time is spent in this document trying to convince the reader that the
+concept is sound. Instead it is simply considered a "solved problem" and as
+such implementing it in XFS is purely an exercise in software engineering.
+
+The fundamental requirements for delayed logging in XFS are simple:
+
+	1. Reduce the amount of metadata written to the log by at least
+	   an order of magnitude.
+	2. Supply sufficient statistics to validate Requirement #1.
+	3. Supply sufficient new tracing infrastructure to be able to debug
+	   problems with the new code.
+	4. No on-disk format change (metadata or log format).
+	5. Enable and disable with a mount option.
+	6. No performance regressions for synchronous transaction workloads.
+
+Delayed Logging: Design
+=======================
+
+Storing Changes
+---------------
+
+The problem with accumulating changes at a logical level (i.e. just using the
+existing log item dirty region tracking) is that when it comes to writing the
+changes to the log buffers, we need to ensure that the object we are formatting
+is not changing while we do this. This requires locking the object to prevent
+concurrent modification. Hence flushing the logical changes to the log would
+require us to lock every object, format them, and then unlock them again.
+
+This introduces lots of scope for deadlocks with transactions that are already
+running. For example, a transaction has object A locked and modified, but needs
+the delayed logging tracking lock to commit the transaction. However, the
+flushing thread has the delayed logging tracking lock already held, and is
+trying to get the lock on object A to flush it to the log buffer. This appears
+to be an unsolvable deadlock condition, and it was solving this problem that
+was the barrier to implementing delayed logging for so long.
+
+The solution is relatively simple - it just took a long time to recognise it.
+Put simply, the current logging code formats the changes to each item into an
+vector array that points to the changed regions in the item. The log write code
+simply copies the memory these vectors point to into the log buffer during
+transaction commit while the item is locked in the transaction. Instead of
+using the log buffer as the destination of the formatting code, we can use an
+allocated memory buffer big enough to fit the formatted vector.
+
+If we then copy the vector into the memory buffer and rewrite the vector to
+point to the memory buffer rather than the object itself, we now have a copy of
+the changes in a format that is compatible with the log buffer writing code.
+that does not require us to lock the item to access. This formatting and
+rewriting can all be done while the object is locked during transaction commit,
+resulting in a vector that is transactionally consistent and can be accessed
+without needing to lock the owning item.
+
+Hence we avoid the need to lock items when we need to flush outstanding
+asynchronous transactions to the log. The differences between the existing
+formatting method and the delayed logging formatting can be seen in the
+diagram below.
+
+Current format log vector::
+
+    Object    +---------------------------------------------+
+    Vector 1      +----+
+    Vector 2                    +----+
+    Vector 3                                   +----------+
+
+After formatting::
+
+    Log Buffer    +-V1-+-V2-+----V3----+
+
+Delayed logging vector::
+
+    Object    +---------------------------------------------+
+    Vector 1      +----+
+    Vector 2                    +----+
+    Vector 3                                   +----------+
+
+After formatting::
+
+    Memory Buffer +-V1-+-V2-+----V3----+
+    Vector 1      +----+
+    Vector 2           +----+
+    Vector 3                +----------+
+
+The memory buffer and associated vector need to be passed as a single object,
+but still need to be associated with the parent object so if the object is
+relogged we can replace the current memory buffer with a new memory buffer that
+contains the latest changes.
+
+The reason for keeping the vector around after we've formatted the memory
+buffer is to support splitting vectors across log buffer boundaries correctly.
+If we don't keep the vector around, we do not know where the region boundaries
+are in the item, so we'd need a new encapsulation method for regions in the log
+buffer writing (i.e. double encapsulation). This would be an on-disk format
+change and as such is not desirable.  It also means we'd have to write the log
+region headers in the formatting stage, which is problematic as there is per
+region state that needs to be placed into the headers during the log write.
+
+Hence we need to keep the vector, but by attaching the memory buffer to it and
+rewriting the vector addresses to point at the memory buffer we end up with a
+self-describing object that can be passed to the log buffer write code to be
+handled in exactly the same manner as the existing log vectors are handled.
+Hence we avoid needing a new on-disk format to handle items that have been
+relogged in memory.
+
+
+Tracking Changes
+----------------
+
+Now that we can record transactional changes in memory in a form that allows
+them to be used without limitations, we need to be able to track and accumulate
+them so that they can be written to the log at some later point in time.  The
+log item is the natural place to store this vector and buffer, and also makes sense
+to be the object that is used to track committed objects as it will always
+exist once the object has been included in a transaction.
+
+The log item is already used to track the log items that have been written to
+the log but not yet written to disk. Such log items are considered "active"
+and as such are stored in the Active Item List (AIL) which is a LSN-ordered
+double linked list. Items are inserted into this list during log buffer IO
+completion, after which they are unpinned and can be written to disk. An object
+that is in the AIL can be relogged, which causes the object to be pinned again
+and then moved forward in the AIL when the log buffer IO completes for that
+transaction.
+
+Essentially, this shows that an item that is in the AIL can still be modified
+and relogged, so any tracking must be separate to the AIL infrastructure. As
+such, we cannot reuse the AIL list pointers for tracking committed items, nor
+can we store state in any field that is protected by the AIL lock. Hence the
+committed item tracking needs its own locks, lists and state fields in the log
+item.
+
+Similar to the AIL, tracking of committed items is done through a new list
+called the Committed Item List (CIL).  The list tracks log items that have been
+committed and have formatted memory buffers attached to them. It tracks objects
+in transaction commit order, so when an object is relogged it is removed from
+its place in the list and re-inserted at the tail. This is entirely arbitrary
+and done to make it easy for debugging - the last items in the list are the
+ones that are most recently modified. Ordering of the CIL is not necessary for
+transactional integrity (as discussed in the next section) so the ordering is
+done for convenience/sanity of the developers.
+
+
+Delayed Logging: Checkpoints
+----------------------------
+
+When we have a log synchronisation event, commonly known as a "log force",
+all the items in the CIL must be written into the log via the log buffers.
+We need to write these items in the order that they exist in the CIL, and they
+need to be written as an atomic transaction. The need for all the objects to be
+written as an atomic transaction comes from the requirements of relogging and
+log replay - all the changes in all the objects in a given transaction must
+either be completely replayed during log recovery, or not replayed at all. If
+a transaction is not replayed because it is not complete in the log, then
+no later transactions should be replayed, either.
+
+To fulfill this requirement, we need to write the entire CIL in a single log
+transaction. Fortunately, the XFS log code has no fixed limit on the size of a
+transaction, nor does the log replay code. The only fundamental limit is that
+the transaction cannot be larger than just under half the size of the log.  The
+reason for this limit is that to find the head and tail of the log, there must
+be at least one complete transaction in the log at any given time. If a
+transaction is larger than half the log, then there is the possibility that a
+crash during the write of a such a transaction could partially overwrite the
+only complete previous transaction in the log. This will result in a recovery
+failure and an inconsistent filesystem and hence we must enforce the maximum
+size of a checkpoint to be slightly less than a half the log.
+
+Apart from this size requirement, a checkpoint transaction looks no different
+to any other transaction - it contains a transaction header, a series of
+formatted log items and a commit record at the tail. From a recovery
+perspective, the checkpoint transaction is also no different - just a lot
+bigger with a lot more items in it. The worst case effect of this is that we
+might need to tune the recovery transaction object hash size.
+
+Because the checkpoint is just another transaction and all the changes to log
+items are stored as log vectors, we can use the existing log buffer writing
+code to write the changes into the log. To do this efficiently, we need to
+minimise the time we hold the CIL locked while writing the checkpoint
+transaction. The current log write code enables us to do this easily with the
+way it separates the writing of the transaction contents (the log vectors) from
+the transaction commit record, but tracking this requires us to have a
+per-checkpoint context that travels through the log write process through to
+checkpoint completion.
+
+Hence a checkpoint has a context that tracks the state of the current
+checkpoint from initiation to checkpoint completion. A new context is initiated
+at the same time a checkpoint transaction is started. That is, when we remove
+all the current items from the CIL during a checkpoint operation, we move all
+those changes into the current checkpoint context. We then initialise a new
+context and attach that to the CIL for aggregation of new transactions.
+
+This allows us to unlock the CIL immediately after transfer of all the
+committed items and effectively allows new transactions to be issued while we
+are formatting the checkpoint into the log. It also allows concurrent
+checkpoints to be written into the log buffers in the case of log force heavy
+workloads, just like the existing transaction commit code does. This, however,
+requires that we strictly order the commit records in the log so that
+checkpoint sequence order is maintained during log replay.
+
+To ensure that we can be writing an item into a checkpoint transaction at
+the same time another transaction modifies the item and inserts the log item
+into the new CIL, then checkpoint transaction commit code cannot use log items
+to store the list of log vectors that need to be written into the transaction.
+Hence log vectors need to be able to be chained together to allow them to be
+detached from the log items. That is, when the CIL is flushed the memory
+buffer and log vector attached to each log item needs to be attached to the
+checkpoint context so that the log item can be released. In diagrammatic form,
+the CIL would look like this before the flush::
+
+	CIL Head
+	   |
+	   V
+	Log Item <-> log vector 1	-> memory buffer
+	   |				-> vector array
+	   V
+	Log Item <-> log vector 2	-> memory buffer
+	   |				-> vector array
+	   V
+	......
+	   |
+	   V
+	Log Item <-> log vector N-1	-> memory buffer
+	   |				-> vector array
+	   V
+	Log Item <-> log vector N	-> memory buffer
+					-> vector array
+
+And after the flush the CIL head is empty, and the checkpoint context log
+vector list would look like::
+
+	Checkpoint Context
+	   |
+	   V
+	log vector 1	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	log vector 2	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	......
+	   |
+	   V
+	log vector N-1	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	log vector N	-> memory buffer
+			-> vector array
+			-> Log Item
+
+Once this transfer is done, the CIL can be unlocked and new transactions can
+start, while the checkpoint flush code works over the log vector chain to
+commit the checkpoint.
+
+Once the checkpoint is written into the log buffers, the checkpoint context is
+attached to the log buffer that the commit record was written to along with a
+completion callback. Log IO completion will call that callback, which can then
+run transaction committed processing for the log items (i.e. insert into AIL
+and unpin) in the log vector chain and then free the log vector chain and
+checkpoint context.
+
+Discussion Point: I am uncertain as to whether the log item is the most
+efficient way to track vectors, even though it seems like the natural way to do
+it. The fact that we walk the log items (in the CIL) just to chain the log
+vectors and break the link between the log item and the log vector means that
+we take a cache line hit for the log item list modification, then another for
+the log vector chaining. If we track by the log vectors, then we only need to
+break the link between the log item and the log vector, which means we should
+dirty only the log item cachelines. Normally I wouldn't be concerned about one
+vs two dirty cachelines except for the fact I've seen upwards of 80,000 log
+vectors in one checkpoint transaction. I'd guess this is a "measure and
+compare" situation that can be done after a working and reviewed implementation
+is in the dev tree....
+
+Delayed Logging: Checkpoint Sequencing
+--------------------------------------
+
+One of the key aspects of the XFS transaction subsystem is that it tags
+committed transactions with the log sequence number of the transaction commit.
+This allows transactions to be issued asynchronously even though there may be
+future operations that cannot be completed until that transaction is fully
+committed to the log. In the rare case that a dependent operation occurs (e.g.
+re-using a freed metadata extent for a data extent), a special, optimised log
+force can be issued to force the dependent transaction to disk immediately.
+
+To do this, transactions need to record the LSN of the commit record of the
+transaction. This LSN comes directly from the log buffer the transaction is
+written into. While this works just fine for the existing transaction
+mechanism, it does not work for delayed logging because transactions are not
+written directly into the log buffers. Hence some other method of sequencing
+transactions is required.
+
+As discussed in the checkpoint section, delayed logging uses per-checkpoint
+contexts, and as such it is simple to assign a sequence number to each
+checkpoint. Because the switching of checkpoint contexts must be done
+atomically, it is simple to ensure that each new context has a monotonically
+increasing sequence number assigned to it without the need for an external
+atomic counter - we can just take the current context sequence number and add
+one to it for the new context.
+
+Then, instead of assigning a log buffer LSN to the transaction commit LSN
+during the commit, we can assign the current checkpoint sequence. This allows
+operations that track transactions that have not yet completed know what
+checkpoint sequence needs to be committed before they can continue. As a
+result, the code that forces the log to a specific LSN now needs to ensure that
+the log forces to a specific checkpoint.
+
+To ensure that we can do this, we need to track all the checkpoint contexts
+that are currently committing to the log. When we flush a checkpoint, the
+context gets added to a "committing" list which can be searched. When a
+checkpoint commit completes, it is removed from the committing list. Because
+the checkpoint context records the LSN of the commit record for the checkpoint,
+we can also wait on the log buffer that contains the commit record, thereby
+using the existing log force mechanisms to execute synchronous forces.
+
+It should be noted that the synchronous forces may need to be extended with
+mitigation algorithms similar to the current log buffer code to allow
+aggregation of multiple synchronous transactions if there are already
+synchronous transactions being flushed. Investigation of the performance of the
+current design is needed before making any decisions here.
+
+The main concern with log forces is to ensure that all the previous checkpoints
+are also committed to disk before the one we need to wait for. Therefore we
+need to check that all the prior contexts in the committing list are also
+complete before waiting on the one we need to complete. We do this
+synchronisation in the log force code so that we don't need to wait anywhere
+else for such serialisation - it only matters when we do a log force.
+
+The only remaining complexity is that a log force now also has to handle the
+case where the forcing sequence number is the same as the current context. That
+is, we need to flush the CIL and potentially wait for it to complete. This is a
+simple addition to the existing log forcing code to check the sequence numbers
+and push if required. Indeed, placing the current sequence checkpoint flush in
+the log force code enables the current mechanism for issuing synchronous
+transactions to remain untouched (i.e. commit an asynchronous transaction, then
+force the log at the LSN of that transaction) and so the higher level code
+behaves the same regardless of whether delayed logging is being used or not.
+
+Delayed Logging: Checkpoint Log Space Accounting
+------------------------------------------------
+
+The big issue for a checkpoint transaction is the log space reservation for the
+transaction. We don't know how big a checkpoint transaction is going to be
+ahead of time, nor how many log buffers it will take to write out, nor the
+number of split log vector regions are going to be used. We can track the
+amount of log space required as we add items to the commit item list, but we
+still need to reserve the space in the log for the checkpoint.
+
+A typical transaction reserves enough space in the log for the worst case space
+usage of the transaction. The reservation accounts for log record headers,
+transaction and region headers, headers for split regions, buffer tail padding,
+etc. as well as the actual space for all the changed metadata in the
+transaction. While some of this is fixed overhead, much of it is dependent on
+the size of the transaction and the number of regions being logged (the number
+of log vectors in the transaction).
+
+An example of the differences would be logging directory changes versus logging
+inode changes. If you modify lots of inode cores (e.g. ``chmod -R g+w *``), then
+there are lots of transactions that only contain an inode core and an inode log
+format structure. That is, two vectors totaling roughly 150 bytes. If we modify
+10,000 inodes, we have about 1.5MB of metadata to write in 20,000 vectors. Each
+vector is 12 bytes, so the total to be logged is approximately 1.75MB. In
+comparison, if we are logging full directory buffers, they are typically 4KB
+each, so we in 1.5MB of directory buffers we'd have roughly 400 buffers and a
+buffer format structure for each buffer - roughly 800 vectors or 1.51MB total
+space.  From this, it should be obvious that a static log space reservation is
+not particularly flexible and is difficult to select the "optimal value" for
+all workloads.
+
+Further, if we are going to use a static reservation, which bit of the entire
+reservation does it cover? We account for space used by the transaction
+reservation by tracking the space currently used by the object in the CIL and
+then calculating the increase or decrease in space used as the object is
+relogged. This allows for a checkpoint reservation to only have to account for
+log buffer metadata used such as log header records.
+
+However, even using a static reservation for just the log metadata is
+problematic. Typically log record headers use at least 16KB of log space per
+1MB of log space consumed (512 bytes per 32k) and the reservation needs to be
+large enough to handle arbitrary sized checkpoint transactions. This
+reservation needs to be made before the checkpoint is started, and we need to
+be able to reserve the space without sleeping.  For a 8MB checkpoint, we need a
+reservation of around 150KB, which is a non-trivial amount of space.
+
+A static reservation needs to manipulate the log grant counters - we can take a
+permanent reservation on the space, but we still need to make sure we refresh
+the write reservation (the actual space available to the transaction) after
+every checkpoint transaction completion. Unfortunately, if this space is not
+available when required, then the regrant code will sleep waiting for it.
+
+The problem with this is that it can lead to deadlocks as we may need to commit
+checkpoints to be able to free up log space (refer back to the description of
+rolling transactions for an example of this).  Hence we *must* always have
+space available in the log if we are to use static reservations, and that is
+very difficult and complex to arrange. It is possible to do, but there is a
+simpler way.
+
+The simpler way of doing this is tracking the entire log space used by the
+items in the CIL and using this to dynamically calculate the amount of log
+space required by the log metadata. If this log metadata space changes as a
+result of a transaction commit inserting a new memory buffer into the CIL, then
+the difference in space required is removed from the transaction that causes
+the change. Transactions at this level will *always* have enough space
+available in their reservation for this as they have already reserved the
+maximal amount of log metadata space they require, and such a delta reservation
+will always be less than or equal to the maximal amount in the reservation.
+
+Hence we can grow the checkpoint transaction reservation dynamically as items
+are added to the CIL and avoid the need for reserving and regranting log space
+up front. This avoids deadlocks and removes a blocking point from the
+checkpoint flush code.
+
+As mentioned early, transactions can't grow to more than half the size of the
+log. Hence as part of the reservation growing, we need to also check the size
+of the reservation against the maximum allowed transaction size. If we reach
+the maximum threshold, we need to push the CIL to the log. This is effectively
+a "background flush" and is done on demand. This is identical to
+a CIL push triggered by a log force, only that there is no waiting for the
+checkpoint commit to complete. This background push is checked and executed by
+transaction commit code.
+
+If the transaction subsystem goes idle while we still have items in the CIL,
+they will be flushed by the periodic log force issued by the xfssyncd. This log
+force will push the CIL to disk, and if the transaction subsystem stays idle,
+allow the idle log to be covered (effectively marked clean) in exactly the same
+manner that is done for the existing logging method. A discussion point is
+whether this log force needs to be done more frequently than the current rate
+which is once every 30s.
+
+
+Delayed Logging: Log Item Pinning
+---------------------------------
+
+Currently log items are pinned during transaction commit while the items are
+still locked. This happens just after the items are formatted, though it could
+be done any time before the items are unlocked. The result of this mechanism is
+that items get pinned once for every transaction that is committed to the log
+buffers. Hence items that are relogged in the log buffers will have a pin count
+for every outstanding transaction they were dirtied in. When each of these
+transactions is completed, they will unpin the item once. As a result, the item
+only becomes unpinned when all the transactions complete and there are no
+pending transactions. Thus the pinning and unpinning of a log item is symmetric
+as there is a 1:1 relationship with transaction commit and log item completion.
+
+For delayed logging, however, we have an asymmetric transaction commit to
+completion relationship. Every time an object is relogged in the CIL it goes
+through the commit process without a corresponding completion being registered.
+That is, we now have a many-to-one relationship between transaction commit and
+log item completion. The result of this is that pinning and unpinning of the
+log items becomes unbalanced if we retain the "pin on transaction commit, unpin
+on transaction completion" model.
+
+To keep pin/unpin symmetry, the algorithm needs to change to a "pin on
+insertion into the CIL, unpin on checkpoint completion". In other words, the
+pinning and unpinning becomes symmetric around a checkpoint context. We have to
+pin the object the first time it is inserted into the CIL - if it is already in
+the CIL during a transaction commit, then we do not pin it again. Because there
+can be multiple outstanding checkpoint contexts, we can still see elevated pin
+counts, but as each checkpoint completes the pin count will retain the correct
+value according to its context.
+
+Just to make matters slightly more complex, this checkpoint level context
+for the pin count means that the pinning of an item must take place under the
+CIL commit/flush lock. If we pin the object outside this lock, we cannot
+guarantee which context the pin count is associated with. This is because of
+the fact pinning the item is dependent on whether the item is present in the
+current CIL or not. If we don't pin the CIL first before we check and pin the
+object, we have a race with CIL being flushed between the check and the pin
+(or not pinning, as the case may be). Hence we must hold the CIL flush/commit
+lock to guarantee that we pin the items correctly.
+
+Delayed Logging: Concurrent Scalability
+---------------------------------------
+
+A fundamental requirement for the CIL is that accesses through transaction
+commits must scale to many concurrent commits. The current transaction commit
+code does not break down even when there are transactions coming from 2048
+processors at once. The current transaction code does not go any faster than if
+there was only one CPU using it, but it does not slow down either.
+
+As a result, the delayed logging transaction commit code needs to be designed
+for concurrency from the ground up. It is obvious that there are serialisation
+points in the design - the three important ones are:
+
+	1. Locking out new transaction commits while flushing the CIL
+	2. Adding items to the CIL and updating item space accounting
+	3. Checkpoint commit ordering
+
+Looking at the transaction commit and CIL flushing interactions, it is clear
+that we have a many-to-one interaction here. That is, the only restriction on
+the number of concurrent transactions that can be trying to commit at once is
+the amount of space available in the log for their reservations. The practical
+limit here is in the order of several hundred concurrent transactions for a
+128MB log, which means that it is generally one per CPU in a machine.
+
+The amount of time a transaction commit needs to hold out a flush is a
+relatively long period of time - the pinning of log items needs to be done
+while we are holding out a CIL flush, so at the moment that means it is held
+across the formatting of the objects into memory buffers (i.e. while memcpy()s
+are in progress). Ultimately a two pass algorithm where the formatting is done
+separately to the pinning of objects could be used to reduce the hold time of
+the transaction commit side.
+
+Because of the number of potential transaction commit side holders, the lock
+really needs to be a sleeping lock - if the CIL flush takes the lock, we do not
+want every other CPU in the machine spinning on the CIL lock. Given that
+flushing the CIL could involve walking a list of tens of thousands of log
+items, it will get held for a significant time and so spin contention is a
+significant concern. Preventing lots of CPUs spinning doing nothing is the
+main reason for choosing a sleeping lock even though nothing in either the
+transaction commit or CIL flush side sleeps with the lock held.
+
+It should also be noted that CIL flushing is also a relatively rare operation
+compared to transaction commit for asynchronous transaction workloads - only
+time will tell if using a read-write semaphore for exclusion will limit
+transaction commit concurrency due to cache line bouncing of the lock on the
+read side.
+
+The second serialisation point is on the transaction commit side where items
+are inserted into the CIL. Because transactions can enter this code
+concurrently, the CIL needs to be protected separately from the above
+commit/flush exclusion. It also needs to be an exclusive lock but it is only
+held for a very short time and so a spin lock is appropriate here. It is
+possible that this lock will become a contention point, but given the short
+hold time once per transaction I think that contention is unlikely.
+
+The final serialisation point is the checkpoint commit record ordering code
+that is run as part of the checkpoint commit and log force sequencing. The code
+path that triggers a CIL flush (i.e. whatever triggers the log force) will enter
+an ordering loop after writing all the log vectors into the log buffers but
+before writing the commit record. This loop walks the list of committing
+checkpoints and needs to block waiting for checkpoints to complete their commit
+record write. As a result it needs a lock and a wait variable. Log force
+sequencing also requires the same lock, list walk, and blocking mechanism to
+ensure completion of checkpoints.
+
+These two sequencing operations can use the mechanism even though the
+events they are waiting for are different. The checkpoint commit record
+sequencing needs to wait until checkpoint contexts contain a commit LSN
+(obtained through completion of a commit record write) while log force
+sequencing needs to wait until previous checkpoint contexts are removed from
+the committing list (i.e. they've completed). A simple wait variable and
+broadcast wakeups (thundering herds) has been used to implement these two
+serialisation queues. They use the same lock as the CIL, too. If we see too
+much contention on the CIL lock, or too many context switches as a result of
+the broadcast wakeups these operations can be put under a new spinlock and
+given separate wait lists to reduce lock contention and the number of processes
+woken by the wrong event.
+
+
+Lifecycle Changes
+-----------------
+
+The existing log item life cycle is as follows::
+
+	1. Transaction allocate
+	2. Transaction reserve
+	3. Lock item
+	4. Join item to transaction
+		If not already attached,
+			Allocate log item
+			Attach log item to owner item
+		Attach log item to transaction
+	5. Modify item
+		Record modifications in log item
+	6. Transaction commit
+		Pin item in memory
+		Format item into log buffer
+		Write commit LSN into transaction
+		Unlock item
+		Attach transaction to log buffer
+
+	<log buffer IO dispatched>
+	<log buffer IO completes>
+
+	7. Transaction completion
+		Mark log item committed
+		Insert log item into AIL
+			Write commit LSN into log item
+		Unpin log item
+	8. AIL traversal
+		Lock item
+		Mark log item clean
+		Flush item to disk
+
+	<item IO completion>
+
+	9. Log item removed from AIL
+		Moves log tail
+		Item unlocked
+
+Essentially, steps 1-6 operate independently from step 7, which is also
+independent of steps 8-9. An item can be locked in steps 1-6 or steps 8-9
+at the same time step 7 is occurring, but only steps 1-6 or 8-9 can occur
+at the same time. If the log item is in the AIL or between steps 6 and 7
+and steps 1-6 are re-entered, then the item is relogged. Only when steps 8-9
+are entered and completed is the object considered clean.
+
+With delayed logging, there are new steps inserted into the life cycle::
+
+	1. Transaction allocate
+	2. Transaction reserve
+	3. Lock item
+	4. Join item to transaction
+		If not already attached,
+			Allocate log item
+			Attach log item to owner item
+		Attach log item to transaction
+	5. Modify item
+		Record modifications in log item
+	6. Transaction commit
+		Pin item in memory if not pinned in CIL
+		Format item into log vector + buffer
+		Attach log vector and buffer to log item
+		Insert log item into CIL
+		Write CIL context sequence into transaction
+		Unlock item
+
+	<next log force>
+
+	7. CIL push
+		lock CIL flush
+		Chain log vectors and buffers together
+		Remove items from CIL
+		unlock CIL flush
+		write log vectors into log
+		sequence commit records
+		attach checkpoint context to log buffer
+
+	<log buffer IO dispatched>
+	<log buffer IO completes>
+
+	8. Checkpoint completion
+		Mark log item committed
+		Insert item into AIL
+			Write commit LSN into log item
+		Unpin log item
+	9. AIL traversal
+		Lock item
+		Mark log item clean
+		Flush item to disk
+	<item IO completion>
+	10. Log item removed from AIL
+		Moves log tail
+		Item unlocked
+
+From this, it can be seen that the only life cycle differences between the two
+logging methods are in the middle of the life cycle - they still have the same
+beginning and end and execution constraints. The only differences are in the
+committing of the log items to the log itself and the completion processing.
+Hence delayed logging should not introduce any constraints on log item
+behaviour, allocation or freeing that don't already exist.
+
+As a result of this zero-impact "insertion" of delayed logging infrastructure
+and the design of the internal structures to avoid on disk format changes, we
+can basically switch between delayed logging and the existing mechanism with a
+mount option. Fundamentally, there is no reason why the log manager would not
+be able to swap methods automatically and transparently depending on load
+characteristics, but this should not be necessary if delayed logging works as
+designed.
diff --git a/Documentation/filesystems/xfs/xfs-maintainer-entry-profile.rst b/Documentation/filesystems/xfs/xfs-maintainer-entry-profile.rst
new file mode 100644
index 000000000..ce4584fb3
--- /dev/null
+++ b/Documentation/filesystems/xfs/xfs-maintainer-entry-profile.rst
@@ -0,0 +1,194 @@
+XFS Maintainer Entry Profile
+============================
+
+Overview
+--------
+XFS is a well known high-performance filesystem in the Linux kernel.
+The aim of this project is to provide and maintain a robust and
+performant filesystem.
+
+Patches are generally merged to the for-next branch of the appropriate
+git repository.
+After a testing period, the for-next branch is merged to the master
+branch.
+
+Kernel code are merged to the xfs-linux tree[0].
+Userspace code are merged to the xfsprogs tree[1].
+Test cases are merged to the xfstests tree[2].
+Ondisk format documentation are merged to the xfs-documentation tree[3].
+
+All patchsets involving XFS *must* be cc'd in their entirety to the mailing
+list linux-xfs@vger.kernel.org.
+
+Roles
+-----
+There are eight key roles in the XFS project.
+A person can take on multiple roles, and a role can be filled by
+multiple people.
+Anyone taking on a role is advised to check in with themselves and
+others on a regular basis about burnout.
+
+- **Outside Contributor**: Anyone who sends a patch but is not involved
+  in the XFS project on a regular basis.
+  These folks are usually people who work on other filesystems or
+  elsewhere in the kernel community.
+
+- **Developer**: Someone who is familiar with the XFS codebase enough to
+  write new code, documentation, and tests.
+
+  Developers can often be found in the IRC channel mentioned by the ``C:``
+  entry in the kernel MAINTAINERS file.
+
+- **Senior Developer**: A developer who is very familiar with at least
+  some part of the XFS codebase and/or other subsystems in the kernel.
+  These people collectively decide the long term goals of the project
+  and nudge the community in that direction.
+  They should help prioritize development and review work for each release
+  cycle.
+
+  Senior developers tend to be more active participants in the IRC channel.
+
+- **Reviewer**: Someone (most likely also a developer) who reads code
+  submissions to decide:
+
+  0. Is the idea behind the contribution sound?
+  1. Does the idea fit the goals of the project?
+  2. Is the contribution designed correctly?
+  3. Is the contribution polished?
+  4. Can the contribution be tested effectively?
+
+  Reviewers should identify themselves with an ``R:`` entry in the kernel
+  and fstests MAINTAINERS files.
+
+- **Testing Lead**: This person is responsible for setting the test
+  coverage goals of the project, negotiating with developers to decide
+  on new tests for new features, and making sure that developers and
+  release managers execute on the testing.
+
+  The testing lead should identify themselves with an ``M:`` entry in
+  the XFS section of the fstests MAINTAINERS file.
+
+- **Bug Triager**: Someone who examines incoming bug reports in just
+  enough detail to identify the person to whom the report should be
+  forwarded.
+
+  The bug triagers should identify themselves with a ``B:`` entry in
+  the kernel MAINTAINERS file.
+
+- **Release Manager**: This person merges reviewed patchsets into an
+  integration branch, tests the result locally, pushes the branch to a
+  public git repository, and sends pull requests further upstream.
+  The release manager is not expected to work on new feature patchsets.
+  If a developer and a reviewer fail to reach a resolution on some point,
+  the release manager must have the ability to intervene to try to drive a
+  resolution.
+
+  The release manager should identify themselves with an ``M:`` entry in
+  the kernel MAINTAINERS file.
+
+- **Community Manager**: This person calls and moderates meetings of as many
+  XFS participants as they can get when mailing list discussions prove
+  insufficient for collective decisionmaking.
+  They may also serve as liaison between managers of the organizations
+  sponsoring work on any part of XFS.
+
+- **LTS Maintainer**: Someone who backports and tests bug fixes from
+  upstream to the LTS kernels.
+  There tend to be six separate LTS trees at any given time.
+
+  The maintainer for a given LTS release should identify themselves with an
+  ``M:`` entry in the MAINTAINERS file for that LTS tree.
+  Unmaintained LTS kernels should be marked with status ``S: Orphan`` in that
+  same file.
+
+Submission Checklist Addendum
+-----------------------------
+Please follow these additional rules when submitting to XFS:
+
+- Patches affecting only the filesystem itself should be based against
+  the latest -rc or the for-next branch.
+  These patches will be merged back to the for-next branch.
+
+- Authors of patches touching other subsystems need to coordinate with
+  the maintainers of XFS and the relevant subsystems to decide how to
+  proceed with a merge.
+
+- Any patchset changing XFS should be cc'd in its entirety to linux-xfs.
+  Do not send partial patchsets; that makes analysis of the broader
+  context of the changes unnecessarily difficult.
+
+- Anyone making kernel changes that have corresponding changes to the
+  userspace utilities should send the userspace changes as separate
+  patchsets immediately after the kernel patchsets.
+
+- Authors of bug fix patches are expected to use fstests[2] to perform
+  an A/B test of the patch to determine that there are no regressions.
+  When possible, a new regression test case should be written for
+  fstests.
+
+- Authors of new feature patchsets must ensure that fstests will have
+  appropriate functional and input corner-case test cases for the new
+  feature.
+
+- When implementing a new feature, it is strongly suggested that the
+  developers write a design document to answer the following questions:
+
+  * **What** problem is this trying to solve?
+
+  * **Who** will benefit from this solution, and **where** will they
+    access it?
+
+  * **How** will this new feature work?  This should touch on major data
+    structures and algorithms supporting the solution at a higher level
+    than code comments.
+
+  * **What** userspace interfaces are necessary to build off of the new
+    features?
+
+  * **How** will this work be tested to ensure that it solves the
+    problems laid out in the design document without causing new
+    problems?
+
+  The design document should be committed in the kernel documentation
+  directory.
+  It may be omitted if the feature is already well known to the
+  community.
+
+- Patchsets for the new tests should be submitted as separate patchsets
+  immediately after the kernel and userspace code patchsets.
+
+- Changes to the on-disk format of XFS must be described in the ondisk
+  format document[3] and submitted as a patchset after the fstests
+  patchsets.
+
+- Patchsets implementing bug fixes and further code cleanups should put
+  the bug fixes at the beginning of the series to ease backporting.
+
+Key Release Cycle Dates
+-----------------------
+Bug fixes may be sent at any time, though the release manager may decide to
+defer a patch when the next merge window is close.
+
+Code submissions targeting the next merge window should be sent between
+-rc1 and -rc6.
+This gives the community time to review the changes, to suggest other changes,
+and for the author to retest those changes.
+
+Code submissions also requiring changes to fs/iomap and targeting the
+next merge window should be sent between -rc1 and -rc4.
+This allows the broader kernel community adequate time to test the
+infrastructure changes.
+
+Review Cadence
+--------------
+In general, please wait at least one week before pinging for feedback.
+To find reviewers, either consult the MAINTAINERS file, or ask
+developers that have Reviewed-by tags for XFS changes to take a look and
+offer their opinion.
+
+References
+----------
+| [0] https://git.kernel.org/pub/scm/fs/xfs/xfs-linux.git/
+| [1] https://git.kernel.org/pub/scm/fs/xfs/xfsprogs-dev.git/
+| [2] https://git.kernel.org/pub/scm/fs/xfs/xfstests-dev.git/
+| [3] https://git.kernel.org/pub/scm/fs/xfs/xfs-documentation.git/
diff --git a/Documentation/filesystems/xfs/xfs-online-fsck-design.rst b/Documentation/filesystems/xfs/xfs-online-fsck-design.rst
new file mode 100644
index 000000000..3d9233f40
--- /dev/null
+++ b/Documentation/filesystems/xfs/xfs-online-fsck-design.rst
@@ -0,0 +1,5279 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. _xfs_online_fsck_design:
+
+..
+        Mapping of heading styles within this document:
+        Heading 1 uses "====" above and below
+        Heading 2 uses "===="
+        Heading 3 uses "----"
+        Heading 4 uses "````"
+        Heading 5 uses "^^^^"
+        Heading 6 uses "~~~~"
+        Heading 7 uses "...."
+
+        Sections are manually numbered because apparently that's what everyone
+        does in the kernel.
+
+======================
+XFS Online Fsck Design
+======================
+
+This document captures the design of the online filesystem check feature for
+XFS.
+The purpose of this document is threefold:
+
+- To help kernel distributors understand exactly what the XFS online fsck
+  feature is, and issues about which they should be aware.
+
+- To help people reading the code to familiarize themselves with the relevant
+  concepts and design points before they start digging into the code.
+
+- To help developers maintaining the system by capturing the reasons
+  supporting higher level decision making.
+
+As the online fsck code is merged, the links in this document to topic branches
+will be replaced with links to code.
+
+This document is licensed under the terms of the GNU Public License, v2.
+The primary author is Darrick J. Wong.
+
+This design document is split into seven parts.
+Part 1 defines what fsck tools are and the motivations for writing a new one.
+Parts 2 and 3 present a high level overview of how online fsck process works
+and how it is tested to ensure correct functionality.
+Part 4 discusses the user interface and the intended usage modes of the new
+program.
+Parts 5 and 6 show off the high level components and how they fit together, and
+then present case studies of how each repair function actually works.
+Part 7 sums up what has been discussed so far and speculates about what else
+might be built atop online fsck.
+
+.. contents:: Table of Contents
+   :local:
+
+1. What is a Filesystem Check?
+==============================
+
+A Unix filesystem has four main responsibilities:
+
+- Provide a hierarchy of names through which application programs can associate
+  arbitrary blobs of data for any length of time,
+
+- Virtualize physical storage media across those names, and
+
+- Retrieve the named data blobs at any time.
+
+- Examine resource usage.
+
+Metadata directly supporting these functions (e.g. files, directories, space
+mappings) are sometimes called primary metadata.
+Secondary metadata (e.g. reverse mapping and directory parent pointers) support
+operations internal to the filesystem, such as internal consistency checking
+and reorganization.
+Summary metadata, as the name implies, condense information contained in
+primary metadata for performance reasons.
+
+The filesystem check (fsck) tool examines all the metadata in a filesystem
+to look for errors.
+In addition to looking for obvious metadata corruptions, fsck also
+cross-references different types of metadata records with each other to look
+for inconsistencies.
+People do not like losing data, so most fsck tools also contains some ability
+to correct any problems found.
+As a word of caution -- the primary goal of most Linux fsck tools is to restore
+the filesystem metadata to a consistent state, not to maximize the data
+recovered.
+That precedent will not be challenged here.
+
+Filesystems of the 20th century generally lacked any redundancy in the ondisk
+format, which means that fsck can only respond to errors by erasing files until
+errors are no longer detected.
+More recent filesystem designs contain enough redundancy in their metadata that
+it is now possible to regenerate data structures when non-catastrophic errors
+occur; this capability aids both strategies.
+
++--------------------------------------------------------------------------+
+| **Note**:                                                                |
++--------------------------------------------------------------------------+
+| System administrators avoid data loss by increasing the number of        |
+| separate storage systems through the creation of backups; and they avoid |
+| downtime by increasing the redundancy of each storage system through the |
+| creation of RAID arrays.                                                 |
+| fsck tools address only the first problem.                               |
++--------------------------------------------------------------------------+
+
+TLDR; Show Me the Code!
+-----------------------
+
+Kernel and userspace code has been fully merged as of October 2025.
+
+Each kernel patchset adding an online repair function will use the same branch
+name across the kernel, xfsprogs, and fstests git repos.
+
+Existing Tools
+--------------
+
+The online fsck tool described here will be the third tool in the history of
+XFS (on Linux) to check and repair filesystems.
+Two programs precede it:
+
+The first program, ``xfs_check``, was created as part of the XFS debugger
+(``xfs_db``) and can only be used with unmounted filesystems.
+It walks all metadata in the filesystem looking for inconsistencies in the
+metadata, though it lacks any ability to repair what it finds.
+Due to its high memory requirements and inability to repair things, this
+program is now deprecated and will not be discussed further.
+
+The second program, ``xfs_repair``, was created to be faster and more robust
+than the first program.
+Like its predecessor, it can only be used with unmounted filesystems.
+It uses extent-based in-memory data structures to reduce memory consumption,
+and tries to schedule readahead IO appropriately to reduce I/O waiting time
+while it scans the metadata of the entire filesystem.
+The most important feature of this tool is its ability to respond to
+inconsistencies in file metadata and directory tree by erasing things as needed
+to eliminate problems.
+Space usage metadata are rebuilt from the observed file metadata.
+
+Problem Statement
+-----------------
+
+The current XFS tools leave several problems unsolved:
+
+1. **User programs** suddenly **lose access** to the filesystem when unexpected
+   shutdowns occur as a result of silent corruptions in the metadata.
+   These occur **unpredictably** and often without warning.
+
+2. **Users** experience a **total loss of service** during the recovery period
+   after an **unexpected shutdown** occurs.
+
+3. **Users** experience a **total loss of service** if the filesystem is taken
+   offline to **look for problems** proactively.
+
+4. **Data owners** cannot **check the integrity** of their stored data without
+   reading all of it.
+   This may expose them to substantial billing costs when a linear media scan
+   performed by the storage system administrator might suffice.
+
+5. **System administrators** cannot **schedule** a maintenance window to deal
+   with corruptions if they **lack the means** to assess filesystem health
+   while the filesystem is online.
+
+6. **Fleet monitoring tools** cannot **automate periodic checks** of filesystem
+   health when doing so requires **manual intervention** and downtime.
+
+7. **Users** can be tricked into **doing things they do not desire** when
+   malicious actors **exploit quirks of Unicode** to place misleading names
+   in directories.
+
+Given this definition of the problems to be solved and the actors who would
+benefit, the proposed solution is a third fsck tool that acts on a running
+filesystem.
+
+This new third program has three components: an in-kernel facility to check
+metadata, an in-kernel facility to repair metadata, and a userspace driver
+program to drive fsck activity on a live filesystem.
+``xfs_scrub`` is the name of the driver program.
+The rest of this document presents the goals and use cases of the new fsck
+tool, describes its major design points in connection to those goals, and
+discusses the similarities and differences with existing tools.
+
++--------------------------------------------------------------------------+
+| **Note**:                                                                |
++--------------------------------------------------------------------------+
+| Throughout this document, the existing offline fsck tool can also be     |
+| referred to by its current name "``xfs_repair``".                        |
+| The userspace driver program for the new online fsck tool can be         |
+| referred to as "``xfs_scrub``".                                          |
+| The kernel portion of online fsck that validates metadata is called      |
+| "online scrub", and portion of the kernel that fixes metadata is called  |
+| "online repair".                                                         |
++--------------------------------------------------------------------------+
+
+The naming hierarchy is broken up into objects known as directories and files
+and the physical space is split into pieces known as allocation groups.
+Sharding enables better performance on highly parallel systems and helps to
+contain the damage when corruptions occur.
+The division of the filesystem into principal objects (allocation groups and
+inodes) means that there are ample opportunities to perform targeted checks and
+repairs on a subset of the filesystem.
+
+While this is going on, other parts continue processing IO requests.
+Even if a piece of filesystem metadata can only be regenerated by scanning the
+entire system, the scan can still be done in the background while other file
+operations continue.
+
+In summary, online fsck takes advantage of resource sharding and redundant
+metadata to enable targeted checking and repair operations while the system
+is running.
+This capability will be coupled to automatic system management so that
+autonomous self-healing of XFS maximizes service availability.
+
+2. Theory of Operation
+======================
+
+Because it is necessary for online fsck to lock and scan live metadata objects,
+online fsck consists of three separate code components.
+The first is the userspace driver program ``xfs_scrub``, which is responsible
+for identifying individual metadata items, scheduling work items for them,
+reacting to the outcomes appropriately, and reporting results to the system
+administrator.
+The second and third are in the kernel, which implements functions to check
+and repair each type of online fsck work item.
+
++------------------------------------------------------------------+
+| **Note**:                                                        |
++------------------------------------------------------------------+
+| For brevity, this document shortens the phrase "online fsck work |
+| item" to "scrub item".                                           |
++------------------------------------------------------------------+
+
+Scrub item types are delineated in a manner consistent with the Unix design
+philosophy, which is to say that each item should handle one aspect of a
+metadata structure, and handle it well.
+
+Scope
+-----
+
+In principle, online fsck should be able to check and to repair everything that
+the offline fsck program can handle.
+However, online fsck cannot be running 100% of the time, which means that
+latent errors may creep in after a scrub completes.
+If these errors cause the next mount to fail, offline fsck is the only
+solution.
+This limitation means that maintenance of the offline fsck tool will continue.
+A second limitation of online fsck is that it must follow the same resource
+sharing and lock acquisition rules as the regular filesystem.
+This means that scrub cannot take *any* shortcuts to save time, because doing
+so could lead to concurrency problems.
+In other words, online fsck is not a complete replacement for offline fsck, and
+a complete run of online fsck may take longer than offline fsck.
+However, both of these limitations are acceptable tradeoffs to satisfy the
+different motivations of online fsck, which are to **minimize system downtime**
+and to **increase predictability of operation**.
+
+.. _scrubphases:
+
+Phases of Work
+--------------
+
+The userspace driver program ``xfs_scrub`` splits the work of checking and
+repairing an entire filesystem into seven phases.
+Each phase concentrates on checking specific types of scrub items and depends
+on the success of all previous phases.
+The seven phases are as follows:
+
+1. Collect geometry information about the mounted filesystem and computer,
+   discover the online fsck capabilities of the kernel, and open the
+   underlying storage devices.
+
+2. Check allocation group metadata, all realtime volume metadata, and all quota
+   files.
+   Each metadata structure is scheduled as a separate scrub item.
+   If corruption is found in the inode header or inode btree and ``xfs_scrub``
+   is permitted to perform repairs, then those scrub items are repaired to
+   prepare for phase 3.
+   Repairs are implemented by using the information in the scrub item to
+   resubmit the kernel scrub call with the repair flag enabled; this is
+   discussed in the next section.
+   Optimizations and all other repairs are deferred to phase 4.
+
+3. Check all metadata of every file in the filesystem.
+   Each metadata structure is also scheduled as a separate scrub item.
+   If repairs are needed and ``xfs_scrub`` is permitted to perform repairs,
+   and there were no problems detected during phase 2, then those scrub items
+   are repaired immediately.
+   Optimizations, deferred repairs, and unsuccessful repairs are deferred to
+   phase 4.
+
+4. All remaining repairs and scheduled optimizations are performed during this
+   phase, if the caller permits them.
+   Before starting repairs, the summary counters are checked and any necessary
+   repairs are performed so that subsequent repairs will not fail the resource
+   reservation step due to wildly incorrect summary counters.
+   Unsuccessful repairs are requeued as long as forward progress on repairs is
+   made somewhere in the filesystem.
+   Free space in the filesystem is trimmed at the end of phase 4 if the
+   filesystem is clean.
+
+5. By the start of this phase, all primary and secondary filesystem metadata
+   must be correct.
+   Summary counters such as the free space counts and quota resource counts
+   are checked and corrected.
+   Directory entry names and extended attribute names are checked for
+   suspicious entries such as control characters or confusing Unicode sequences
+   appearing in names.
+
+6. If the caller asks for a media scan, read all allocated and written data
+   file extents in the filesystem.
+   The ability to use hardware-assisted data file integrity checking is new
+   to online fsck; neither of the previous tools have this capability.
+   If media errors occur, they will be mapped to the owning files and reported.
+
+7. Re-check the summary counters and presents the caller with a summary of
+   space usage and file counts.
+
+This allocation of responsibilities will be :ref:`revisited <scrubcheck>`
+later in this document.
+
+Steps for Each Scrub Item
+-------------------------
+
+The kernel scrub code uses a three-step strategy for checking and repairing
+the one aspect of a metadata object represented by a scrub item:
+
+1. The scrub item of interest is checked for corruptions; opportunities for
+   optimization; and for values that are directly controlled by the system
+   administrator but look suspicious.
+   If the item is not corrupt or does not need optimization, resource are
+   released and the positive scan results are returned to userspace.
+   If the item is corrupt or could be optimized but the caller does not permit
+   this, resources are released and the negative scan results are returned to
+   userspace.
+   Otherwise, the kernel moves on to the second step.
+
+2. The repair function is called to rebuild the data structure.
+   Repair functions generally choose rebuild a structure from other metadata
+   rather than try to salvage the existing structure.
+   If the repair fails, the scan results from the first step are returned to
+   userspace.
+   Otherwise, the kernel moves on to the third step.
+
+3. In the third step, the kernel runs the same checks over the new metadata
+   item to assess the efficacy of the repairs.
+   The results of the reassessment are returned to userspace.
+
+Classification of Metadata
+--------------------------
+
+Each type of metadata object (and therefore each type of scrub item) is
+classified as follows:
+
+Primary Metadata
+````````````````
+
+Metadata structures in this category should be most familiar to filesystem
+users either because they are directly created by the user or they index
+objects created by the user
+Most filesystem objects fall into this class:
+
+- Free space and reference count information
+
+- Inode records and indexes
+
+- Storage mapping information for file data
+
+- Directories
+
+- Extended attributes
+
+- Symbolic links
+
+- Quota limits
+
+Scrub obeys the same rules as regular filesystem accesses for resource and lock
+acquisition.
+
+Primary metadata objects are the simplest for scrub to process.
+The principal filesystem object (either an allocation group or an inode) that
+owns the item being scrubbed is locked to guard against concurrent updates.
+The check function examines every record associated with the type for obvious
+errors and cross-references healthy records against other metadata to look for
+inconsistencies.
+Repairs for this class of scrub item are simple, since the repair function
+starts by holding all the resources acquired in the previous step.
+The repair function scans available metadata as needed to record all the
+observations needed to complete the structure.
+Next, it stages the observations in a new ondisk structure and commits it
+atomically to complete the repair.
+Finally, the storage from the old data structure are carefully reaped.
+
+Because ``xfs_scrub`` locks a primary object for the duration of the repair,
+this is effectively an offline repair operation performed on a subset of the
+filesystem.
+This minimizes the complexity of the repair code because it is not necessary to
+handle concurrent updates from other threads, nor is it necessary to access
+any other part of the filesystem.
+As a result, indexed structures can be rebuilt very quickly, and programs
+trying to access the damaged structure will be blocked until repairs complete.
+The only infrastructure needed by the repair code are the staging area for
+observations and a means to write new structures to disk.
+Despite these limitations, the advantage that online repair holds is clear:
+targeted work on individual shards of the filesystem avoids total loss of
+service.
+
+This mechanism is described in section 2.1 ("Off-Line Algorithm") of
+V. Srinivasan and M. J. Carey, `"Performance of On-Line Index Construction
+Algorithms" <https://minds.wisconsin.edu/bitstream/handle/1793/59524/TR1047.pdf>`_,
+*Extending Database Technology*, pp. 293-309, 1992.
+
+Most primary metadata repair functions stage their intermediate results in an
+in-memory array prior to formatting the new ondisk structure, which is very
+similar to the list-based algorithm discussed in section 2.3 ("List-Based
+Algorithms") of Srinivasan.
+However, any data structure builder that maintains a resource lock for the
+duration of the repair is *always* an offline algorithm.
+
+.. _secondary_metadata:
+
+Secondary Metadata
+``````````````````
+
+Metadata structures in this category reflect records found in primary metadata,
+but are only needed for online fsck or for reorganization of the filesystem.
+
+Secondary metadata include:
+
+- Reverse mapping information
+
+- Directory parent pointers
+
+This class of metadata is difficult for scrub to process because scrub attaches
+to the secondary object but needs to check primary metadata, which runs counter
+to the usual order of resource acquisition.
+Frequently, this means that full filesystems scans are necessary to rebuild the
+metadata.
+Check functions can be limited in scope to reduce runtime.
+Repairs, however, require a full scan of primary metadata, which can take a
+long time to complete.
+Under these conditions, ``xfs_scrub`` cannot lock resources for the entire
+duration of the repair.
+
+Instead, repair functions set up an in-memory staging structure to store
+observations.
+Depending on the requirements of the specific repair function, the staging
+index will either have the same format as the ondisk structure or a design
+specific to that repair function.
+The next step is to release all locks and start the filesystem scan.
+When the repair scanner needs to record an observation, the staging data are
+locked long enough to apply the update.
+While the filesystem scan is in progress, the repair function hooks the
+filesystem so that it can apply pending filesystem updates to the staging
+information.
+Once the scan is done, the owning object is re-locked, the live data is used to
+write a new ondisk structure, and the repairs are committed atomically.
+The hooks are disabled and the staging area is freed.
+Finally, the storage from the old data structure are carefully reaped.
+
+Introducing concurrency helps online repair avoid various locking problems, but
+comes at a high cost to code complexity.
+Live filesystem code has to be hooked so that the repair function can observe
+updates in progress.
+The staging area has to become a fully functional parallel structure so that
+updates can be merged from the hooks.
+Finally, the hook, the filesystem scan, and the inode locking model must be
+sufficiently well integrated that a hook event can decide if a given update
+should be applied to the staging structure.
+
+In theory, the scrub implementation could apply these same techniques for
+primary metadata, but doing so would make it massively more complex and less
+performant.
+Programs attempting to access the damaged structures are not blocked from
+operation, which may cause application failure or an unplanned filesystem
+shutdown.
+
+Inspiration for the secondary metadata repair strategy was drawn from section
+2.4 of Srinivasan above, and sections 2 ("NSF: Index Build Without Side-File")
+and 3.1.1 ("Duplicate Key Insert Problem") in C. Mohan, `"Algorithms for
+Creating Indexes for Very Large Tables Without Quiescing Updates"
+<https://dl.acm.org/doi/10.1145/130283.130337>`_, 1992.
+
+The sidecar index mentioned above bears some resemblance to the side file
+method mentioned in Srinivasan and Mohan.
+Their method consists of an index builder that extracts relevant record data to
+build the new structure as quickly as possible; and an auxiliary structure that
+captures all updates that would be committed to the index by other threads were
+the new index already online.
+After the index building scan finishes, the updates recorded in the side file
+are applied to the new index.
+To avoid conflicts between the index builder and other writer threads, the
+builder maintains a publicly visible cursor that tracks the progress of the
+scan through the record space.
+To avoid duplication of work between the side file and the index builder, side
+file updates are elided when the record ID for the update is greater than the
+cursor position within the record ID space.
+
+To minimize changes to the rest of the codebase, XFS online repair keeps the
+replacement index hidden until it's completely ready to go.
+In other words, there is no attempt to expose the keyspace of the new index
+while repair is running.
+The complexity of such an approach would be very high and perhaps more
+appropriate to building *new* indices.
+
+**Future Work Question**: Can the full scan and live update code used to
+facilitate a repair also be used to implement a comprehensive check?
+
+*Answer*: In theory, yes.  Check would be much stronger if each scrub function
+employed these live scans to build a shadow copy of the metadata and then
+compared the shadow records to the ondisk records.
+However, doing that is a fair amount more work than what the checking functions
+do now.
+The live scans and hooks were developed much later.
+That in turn increases the runtime of those scrub functions.
+
+Summary Information
+```````````````````
+
+Metadata structures in this last category summarize the contents of primary
+metadata records.
+These are often used to speed up resource usage queries, and are many times
+smaller than the primary metadata which they represent.
+
+Examples of summary information include:
+
+- Summary counts of free space and inodes
+
+- File link counts from directories
+
+- Quota resource usage counts
+
+Check and repair require full filesystem scans, but resource and lock
+acquisition follow the same paths as regular filesystem accesses.
+
+The superblock summary counters have special requirements due to the underlying
+implementation of the incore counters, and will be treated separately.
+Check and repair of the other types of summary counters (quota resource counts
+and file link counts) employ the same filesystem scanning and hooking
+techniques as outlined above, but because the underlying data are sets of
+integer counters, the staging data need not be a fully functional mirror of the
+ondisk structure.
+
+Inspiration for quota and file link count repair strategies were drawn from
+sections 2.12 ("Online Index Operations") through 2.14 ("Incremental View
+Maintenance") of G.  Graefe, `"Concurrent Queries and Updates in Summary Views
+and Their Indexes"
+<http://www.odbms.org/wp-content/uploads/2014/06/Increment-locks.pdf>`_, 2011.
+
+Since quotas are non-negative integer counts of resource usage, online
+quotacheck can use the incremental view deltas described in section 2.14 to
+track pending changes to the block and inode usage counts in each transaction,
+and commit those changes to a dquot side file when the transaction commits.
+Delta tracking is necessary for dquots because the index builder scans inodes,
+whereas the data structure being rebuilt is an index of dquots.
+Link count checking combines the view deltas and commit step into one because
+it sets attributes of the objects being scanned instead of writing them to a
+separate data structure.
+Each online fsck function will be discussed as case studies later in this
+document.
+
+Risk Management
+---------------
+
+During the development of online fsck, several risk factors were identified
+that may make the feature unsuitable for certain distributors and users.
+Steps can be taken to mitigate or eliminate those risks, though at a cost to
+functionality.
+
+- **Decreased performance**: Adding metadata indices to the filesystem
+  increases the time cost of persisting changes to disk, and the reverse space
+  mapping and directory parent pointers are no exception.
+  System administrators who require the maximum performance can disable the
+  reverse mapping features at format time, though this choice dramatically
+  reduces the ability of online fsck to find inconsistencies and repair them.
+
+- **Incorrect repairs**: As with all software, there might be defects in the
+  software that result in incorrect repairs being written to the filesystem.
+  Systematic fuzz testing (detailed in the next section) is employed by the
+  authors to find bugs early, but it might not catch everything.
+  The kernel build system provides Kconfig options (``CONFIG_XFS_ONLINE_SCRUB``
+  and ``CONFIG_XFS_ONLINE_REPAIR``) to enable distributors to choose not to
+  accept this risk.
+  The xfsprogs build system has a configure option (``--enable-scrub=no``) that
+  disables building of the ``xfs_scrub`` binary, though this is not a risk
+  mitigation if the kernel functionality remains enabled.
+
+- **Inability to repair**: Sometimes, a filesystem is too badly damaged to be
+  repairable.
+  If the keyspaces of several metadata indices overlap in some manner but a
+  coherent narrative cannot be formed from records collected, then the repair
+  fails.
+  To reduce the chance that a repair will fail with a dirty transaction and
+  render the filesystem unusable, the online repair functions have been
+  designed to stage and validate all new records before committing the new
+  structure.
+
+- **Misbehavior**: Online fsck requires many privileges -- raw IO to block
+  devices, opening files by handle, ignoring Unix discretionary access control,
+  and the ability to perform administrative changes.
+  Running this automatically in the background scares people, so the systemd
+  background service is configured to run with only the privileges required.
+  Obviously, this cannot address certain problems like the kernel crashing or
+  deadlocking, but it should be sufficient to prevent the scrub process from
+  escaping and reconfiguring the system.
+  The cron job does not have this protection.
+
+- **Fuzz Kiddiez**: There are many people now who seem to think that running
+  automated fuzz testing of ondisk artifacts to find mischievous behavior and
+  spraying exploit code onto the public mailing list for instant zero-day
+  disclosure is somehow of some social benefit.
+  In the view of this author, the benefit is realized only when the fuzz
+  operators help to **fix** the flaws, but this opinion apparently is not
+  widely shared among security "researchers".
+  The XFS maintainers' continuing ability to manage these events presents an
+  ongoing risk to the stability of the development process.
+  Automated testing should front-load some of the risk while the feature is
+  considered EXPERIMENTAL.
+
+Many of these risks are inherent to software programming.
+Despite this, it is hoped that this new functionality will prove useful in
+reducing unexpected downtime.
+
+3. Testing Plan
+===============
+
+As stated before, fsck tools have three main goals:
+
+1. Detect inconsistencies in the metadata;
+
+2. Eliminate those inconsistencies; and
+
+3. Minimize further loss of data.
+
+Demonstrations of correct operation are necessary to build users' confidence
+that the software behaves within expectations.
+Unfortunately, it was not really feasible to perform regular exhaustive testing
+of every aspect of a fsck tool until the introduction of low-cost virtual
+machines with high-IOPS storage.
+With ample hardware availability in mind, the testing strategy for the online
+fsck project involves differential analysis against the existing fsck tools and
+systematic testing of every attribute of every type of metadata object.
+Testing can be split into four major categories, as discussed below.
+
+Integrated Testing with fstests
+-------------------------------
+
+The primary goal of any free software QA effort is to make testing as
+inexpensive and widespread as possible to maximize the scaling advantages of
+community.
+In other words, testing should maximize the breadth of filesystem configuration
+scenarios and hardware setups.
+This improves code quality by enabling the authors of online fsck to find and
+fix bugs early, and helps developers of new features to find integration
+issues earlier in their development effort.
+
+The Linux filesystem community shares a common QA testing suite,
+`fstests <https://git.kernel.org/pub/scm/fs/xfs/xfstests-dev.git/>`_, for
+functional and regression testing.
+Even before development work began on online fsck, fstests (when run on XFS)
+would run both the ``xfs_check`` and ``xfs_repair -n`` commands on the test and
+scratch filesystems between each test.
+This provides a level of assurance that the kernel and the fsck tools stay in
+alignment about what constitutes consistent metadata.
+During development of the online checking code, fstests was modified to run
+``xfs_scrub -n`` between each test to ensure that the new checking code
+produces the same results as the two existing fsck tools.
+
+To start development of online repair, fstests was modified to run
+``xfs_repair`` to rebuild the filesystem's metadata indices between tests.
+This ensures that offline repair does not crash, leave a corrupt filesystem
+after it exists, or trigger complaints from the online check.
+This also established a baseline for what can and cannot be repaired offline.
+To complete the first phase of development of online repair, fstests was
+modified to be able to run ``xfs_scrub`` in a "force rebuild" mode.
+This enables a comparison of the effectiveness of online repair as compared to
+the existing offline repair tools.
+
+General Fuzz Testing of Metadata Blocks
+---------------------------------------
+
+XFS benefits greatly from having a very robust debugging tool, ``xfs_db``.
+
+Before development of online fsck even began, a set of fstests were created
+to test the rather common fault that entire metadata blocks get corrupted.
+This required the creation of fstests library code that can create a filesystem
+containing every possible type of metadata object.
+Next, individual test cases were created to create a test filesystem, identify
+a single block of a specific type of metadata object, trash it with the
+existing ``blocktrash`` command in ``xfs_db``, and test the reaction of a
+particular metadata validation strategy.
+
+This earlier test suite enabled XFS developers to test the ability of the
+in-kernel validation functions and the ability of the offline fsck tool to
+detect and eliminate the inconsistent metadata.
+This part of the test suite was extended to cover online fsck in exactly the
+same manner.
+
+In other words, for a given fstests filesystem configuration:
+
+* For each metadata object existing on the filesystem:
+
+  * Write garbage to it
+
+  * Test the reactions of:
+
+    1. The kernel verifiers to stop obviously bad metadata
+    2. Offline repair (``xfs_repair``) to detect and fix
+    3. Online repair (``xfs_scrub``) to detect and fix
+
+Targeted Fuzz Testing of Metadata Records
+-----------------------------------------
+
+The testing plan for online fsck includes extending the existing fs testing
+infrastructure to provide a much more powerful facility: targeted fuzz testing
+of every metadata field of every metadata object in the filesystem.
+``xfs_db`` can modify every field of every metadata structure in every
+block in the filesystem to simulate the effects of memory corruption and
+software bugs.
+Given that fstests already contains the ability to create a filesystem
+containing every metadata format known to the filesystem, ``xfs_db`` can be
+used to perform exhaustive fuzz testing!
+
+For a given fstests filesystem configuration:
+
+* For each metadata object existing on the filesystem...
+
+  * For each record inside that metadata object...
+
+    * For each field inside that record...
+
+      * For each conceivable type of transformation that can be applied to a bit field...
+
+        1. Clear all bits
+        2. Set all bits
+        3. Toggle the most significant bit
+        4. Toggle the middle bit
+        5. Toggle the least significant bit
+        6. Add a small quantity
+        7. Subtract a small quantity
+        8. Randomize the contents
+
+        * ...test the reactions of:
+
+          1. The kernel verifiers to stop obviously bad metadata
+          2. Offline checking (``xfs_repair -n``)
+          3. Offline repair (``xfs_repair``)
+          4. Online checking (``xfs_scrub -n``)
+          5. Online repair (``xfs_scrub``)
+          6. Both repair tools (``xfs_scrub`` and then ``xfs_repair`` if online repair doesn't succeed)
+
+This is quite the combinatoric explosion!
+
+Fortunately, having this much test coverage makes it easy for XFS developers to
+check the responses of XFS' fsck tools.
+Since the introduction of the fuzz testing framework, these tests have been
+used to discover incorrect repair code and missing functionality for entire
+classes of metadata objects in ``xfs_repair``.
+The enhanced testing was used to finalize the deprecation of ``xfs_check`` by
+confirming that ``xfs_repair`` could detect at least as many corruptions as
+the older tool.
+
+These tests have been very valuable for ``xfs_scrub`` in the same ways -- they
+allow the online fsck developers to compare online fsck against offline fsck,
+and they enable XFS developers to find deficiencies in the code base.
+
+Proposed patchsets include
+`fuzzing baselines
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfstests-dev.git/log/?h=fuzz-baseline>`_.
+
+Stress Testing
+--------------
+
+A unique requirement to online fsck is the ability to operate on a filesystem
+concurrently with regular workloads.
+Although it is of course impossible to run ``xfs_scrub`` with *zero* observable
+impact on the running system, the online repair code should never introduce
+inconsistencies into the filesystem metadata, and regular workloads should
+never notice resource starvation.
+To verify that these conditions are being met, fstests has been enhanced in
+the following ways:
+
+* For each scrub item type, create a test to exercise checking that item type
+  while running ``fsstress``.
+* For each scrub item type, create a test to exercise repairing that item type
+  while running ``fsstress``.
+* Race ``fsstress`` and ``xfs_scrub -n`` to ensure that checking the whole
+  filesystem doesn't cause problems.
+* Race ``fsstress`` and ``xfs_scrub`` in force-rebuild mode to ensure that
+  force-repairing the whole filesystem doesn't cause problems.
+* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
+  freezing and thawing the filesystem.
+* Race ``xfs_scrub`` in check and force-repair mode against ``fsstress`` while
+  remounting the filesystem read-only and read-write.
+* The same, but running ``fsx`` instead of ``fsstress``.  (Not done yet?)
+
+Success is defined by the ability to run all of these tests without observing
+any unexpected filesystem shutdowns due to corrupted metadata, kernel hang
+check warnings, or any other sort of mischief.
+
+4. User Interface
+=================
+
+The primary user of online fsck is the system administrator, just like offline
+repair.
+Online fsck presents two modes of operation to administrators:
+A foreground CLI process for online fsck on demand, and a background service
+that performs autonomous checking and repair.
+
+Checking on Demand
+------------------
+
+For administrators who want the absolute freshest information about the
+metadata in a filesystem, ``xfs_scrub`` can be run as a foreground process on
+a command line.
+The program checks every piece of metadata in the filesystem while the
+administrator waits for the results to be reported, just like the existing
+``xfs_repair`` tool.
+Both tools share a ``-n`` option to perform a read-only scan, and a ``-v``
+option to increase the verbosity of the information reported.
+
+A new feature of ``xfs_scrub`` is the ``-x`` option, which employs the error
+correction capabilities of the hardware to check data file contents.
+The media scan is not enabled by default because it may dramatically increase
+program runtime and consume a lot of bandwidth on older storage hardware.
+
+The output of a foreground invocation is captured in the system log.
+
+The ``xfs_scrub_all`` program walks the list of mounted filesystems and
+initiates ``xfs_scrub`` for each of them in parallel.
+It serializes scans for any filesystems that resolve to the same top level
+kernel block device to prevent resource overconsumption.
+
+Background Service
+------------------
+
+To reduce the workload of system administrators, the ``xfs_scrub`` package
+provides a suite of `systemd <https://systemd.io/>`_ timers and services that
+run online fsck automatically on weekends by default.
+The background service configures scrub to run with as little privilege as
+possible, the lowest CPU and IO priority, and in a CPU-constrained single
+threaded mode.
+This can be tuned by the systemd administrator at any time to suit the latency
+and throughput requirements of customer workloads.
+
+The output of the background service is also captured in the system log.
+If desired, reports of failures (either due to inconsistencies or mere runtime
+errors) can be emailed automatically by setting the ``EMAIL_ADDR`` environment
+variable in the following service files:
+
+* ``xfs_scrub_fail@.service``
+* ``xfs_scrub_media_fail@.service``
+* ``xfs_scrub_all_fail.service``
+
+The decision to enable the background scan is left to the system administrator.
+This can be done by enabling either of the following services:
+
+* ``xfs_scrub_all.timer`` on systemd systems
+* ``xfs_scrub_all.cron`` on non-systemd systems
+
+This automatic weekly scan is configured out of the box to perform an
+additional media scan of all file data once per month.
+This is less foolproof than, say, storing file data block checksums, but much
+more performant if application software provides its own integrity checking,
+redundancy can be provided elsewhere above the filesystem, or the storage
+device's integrity guarantees are deemed sufficient.
+
+The systemd unit file definitions have been subjected to a security audit
+(as of systemd 249) to ensure that the xfs_scrub processes have as little
+access to the rest of the system as possible.
+This was performed via ``systemd-analyze security``, after which privileges
+were restricted to the minimum required, sandboxing was set up to the maximal
+extent possible with sandboxing and system call filtering; and access to the
+filesystem tree was restricted to the minimum needed to start the program and
+access the filesystem being scanned.
+The service definition files restrict CPU usage to 80% of one CPU core, and
+apply as nice of a priority to IO and CPU scheduling as possible.
+This measure was taken to minimize delays in the rest of the filesystem.
+No such hardening has been performed for the cron job.
+
+Health Reporting
+----------------
+
+XFS caches a summary of each filesystem's health status in memory.
+The information is updated whenever ``xfs_scrub`` is run, or whenever
+inconsistencies are detected in the filesystem metadata during regular
+operations.
+System administrators should use the ``health`` command of ``xfs_spaceman`` to
+download this information into a human-readable format.
+If problems have been observed, the administrator can schedule a reduced
+service window to run the online repair tool to correct the problem.
+Failing that, the administrator can decide to schedule a maintenance window to
+run the traditional offline repair tool to correct the problem.
+
+**Future Work Question**: Should the health reporting integrate with the new
+inotify fs error notification system?
+Would it be helpful for sysadmins to have a daemon to listen for corruption
+notifications and initiate a repair?
+
+*Answer*: These questions remain unanswered, but should be a part of the
+conversation with early adopters and potential downstream users of XFS.
+
+5. Kernel Algorithms and Data Structures
+========================================
+
+This section discusses the key algorithms and data structures of the kernel
+code that provide the ability to check and repair metadata while the system
+is running.
+The first chapters in this section reveal the pieces that provide the
+foundation for checking metadata.
+The remainder of this section presents the mechanisms through which XFS
+regenerates itself.
+
+Self Describing Metadata
+------------------------
+
+Starting with XFS version 5 in 2012, XFS updated the format of nearly every
+ondisk block header to record a magic number, a checksum, a universally
+"unique" identifier (UUID), an owner code, the ondisk address of the block,
+and a log sequence number.
+When loading a block buffer from disk, the magic number, UUID, owner, and
+ondisk address confirm that the retrieved block matches the specific owner of
+the current filesystem, and that the information contained in the block is
+supposed to be found at the ondisk address.
+The first three components enable checking tools to disregard alleged metadata
+that doesn't belong to the filesystem, and the fourth component enables the
+filesystem to detect lost writes.
+
+Whenever a file system operation modifies a block, the change is submitted
+to the log as part of a transaction.
+The log then processes these transactions marking them done once they are
+safely persisted to storage.
+The logging code maintains the checksum and the log sequence number of the last
+transactional update.
+Checksums are useful for detecting torn writes and other discrepancies that can
+be introduced between the computer and its storage devices.
+Sequence number tracking enables log recovery to avoid applying out of date
+log updates to the filesystem.
+
+These two features improve overall runtime resiliency by providing a means for
+the filesystem to detect obvious corruption when reading metadata blocks from
+disk, but these buffer verifiers cannot provide any consistency checking
+between metadata structures.
+
+For more information, please see the documentation for
+Documentation/filesystems/xfs/xfs-self-describing-metadata.rst
+
+Reverse Mapping
+---------------
+
+The original design of XFS (circa 1993) is an improvement upon 1980s Unix
+filesystem design.
+In those days, storage density was expensive, CPU time was scarce, and
+excessive seek time could kill performance.
+For performance reasons, filesystem authors were reluctant to add redundancy to
+the filesystem, even at the cost of data integrity.
+Filesystems designers in the early 21st century choose different strategies to
+increase internal redundancy -- either storing nearly identical copies of
+metadata, or more space-efficient encoding techniques.
+
+For XFS, a different redundancy strategy was chosen to modernize the design:
+a secondary space usage index that maps allocated disk extents back to their
+owners.
+By adding a new index, the filesystem retains most of its ability to scale
+well to heavily threaded workloads involving large datasets, since the primary
+file metadata (the directory tree, the file block map, and the allocation
+groups) remain unchanged.
+Like any system that improves redundancy, the reverse-mapping feature increases
+overhead costs for space mapping activities.
+However, it has two critical advantages: first, the reverse index is key to
+enabling online fsck and other requested functionality such as free space
+defragmentation, better media failure reporting, and filesystem shrinking.
+Second, the different ondisk storage format of the reverse mapping btree
+defeats device-level deduplication because the filesystem requires real
+redundancy.
+
++--------------------------------------------------------------------------+
+| **Sidebar**:                                                             |
++--------------------------------------------------------------------------+
+| A criticism of adding the secondary index is that it does nothing to     |
+| improve the robustness of user data storage itself.                      |
+| This is a valid point, but adding a new index for file data block        |
+| checksums increases write amplification by turning data overwrites into  |
+| copy-writes, which age the filesystem prematurely.                       |
+| In keeping with thirty years of precedent, users who want file data      |
+| integrity can supply as powerful a solution as they require.             |
+| As for metadata, the complexity of adding a new secondary index of space |
+| usage is much less than adding volume management and storage device      |
+| mirroring to XFS itself.                                                 |
+| Perfection of RAID and volume management are best left to existing       |
+| layers in the kernel.                                                    |
++--------------------------------------------------------------------------+
+
+The information captured in a reverse space mapping record is as follows:
+
+.. code-block:: c
+
+	struct xfs_rmap_irec {
+	    xfs_agblock_t    rm_startblock;   /* extent start block */
+	    xfs_extlen_t     rm_blockcount;   /* extent length */
+	    uint64_t         rm_owner;        /* extent owner */
+	    uint64_t         rm_offset;       /* offset within the owner */
+	    unsigned int     rm_flags;        /* state flags */
+	};
+
+The first two fields capture the location and size of the physical space,
+in units of filesystem blocks.
+The owner field tells scrub which metadata structure or file inode have been
+assigned this space.
+For space allocated to files, the offset field tells scrub where the space was
+mapped within the file fork.
+Finally, the flags field provides extra information about the space usage --
+is this an attribute fork extent?  A file mapping btree extent?  Or an
+unwritten data extent?
+
+Online filesystem checking judges the consistency of each primary metadata
+record by comparing its information against all other space indices.
+The reverse mapping index plays a key role in the consistency checking process
+because it contains a centralized alternate copy of all space allocation
+information.
+Program runtime and ease of resource acquisition are the only real limits to
+what online checking can consult.
+For example, a file data extent mapping can be checked against:
+
+* The absence of an entry in the free space information.
+* The absence of an entry in the inode index.
+* The absence of an entry in the reference count data if the file is not
+  marked as having shared extents.
+* The correspondence of an entry in the reverse mapping information.
+
+There are several observations to make about reverse mapping indices:
+
+1. Reverse mappings can provide a positive affirmation of correctness if any of
+   the above primary metadata are in doubt.
+   The checking code for most primary metadata follows a path similar to the
+   one outlined above.
+
+2. Proving the consistency of secondary metadata with the primary metadata is
+   difficult because that requires a full scan of all primary space metadata,
+   which is very time intensive.
+   For example, checking a reverse mapping record for a file extent mapping
+   btree block requires locking the file and searching the entire btree to
+   confirm the block.
+   Instead, scrub relies on rigorous cross-referencing during the primary space
+   mapping structure checks.
+
+3. Consistency scans must use non-blocking lock acquisition primitives if the
+   required locking order is not the same order used by regular filesystem
+   operations.
+   For example, if the filesystem normally takes a file ILOCK before taking
+   the AGF buffer lock but scrub wants to take a file ILOCK while holding
+   an AGF buffer lock, scrub cannot block on that second acquisition.
+   This means that forward progress during this part of a scan of the reverse
+   mapping data cannot be guaranteed if system load is heavy.
+
+In summary, reverse mappings play a key role in reconstruction of primary
+metadata.
+The details of how these records are staged, written to disk, and committed
+into the filesystem are covered in subsequent sections.
+
+Checking and Cross-Referencing
+------------------------------
+
+The first step of checking a metadata structure is to examine every record
+contained within the structure and its relationship with the rest of the
+system.
+XFS contains multiple layers of checking to try to prevent inconsistent
+metadata from wreaking havoc on the system.
+Each of these layers contributes information that helps the kernel to make
+three decisions about the health of a metadata structure:
+
+- Is a part of this structure obviously corrupt (``XFS_SCRUB_OFLAG_CORRUPT``) ?
+- Is this structure inconsistent with the rest of the system
+  (``XFS_SCRUB_OFLAG_XCORRUPT``) ?
+- Is there so much damage around the filesystem that cross-referencing is not
+  possible (``XFS_SCRUB_OFLAG_XFAIL``) ?
+- Can the structure be optimized to improve performance or reduce the size of
+  metadata (``XFS_SCRUB_OFLAG_PREEN``) ?
+- Does the structure contain data that is not inconsistent but deserves review
+  by the system administrator (``XFS_SCRUB_OFLAG_WARNING``) ?
+
+The following sections describe how the metadata scrubbing process works.
+
+Metadata Buffer Verification
+````````````````````````````
+
+The lowest layer of metadata protection in XFS are the metadata verifiers built
+into the buffer cache.
+These functions perform inexpensive internal consistency checking of the block
+itself, and answer these questions:
+
+- Does the block belong to this filesystem?
+
+- Does the block belong to the structure that asked for the read?
+  This assumes that metadata blocks only have one owner, which is always true
+  in XFS.
+
+- Is the type of data stored in the block within a reasonable range of what
+  scrub is expecting?
+
+- Does the physical location of the block match the location it was read from?
+
+- Does the block checksum match the data?
+
+The scope of the protections here are very limited -- verifiers can only
+establish that the filesystem code is reasonably free of gross corruption bugs
+and that the storage system is reasonably competent at retrieval.
+Corruption problems observed at runtime cause the generation of health reports,
+failed system calls, and in the extreme case, filesystem shutdowns if the
+corrupt metadata force the cancellation of a dirty transaction.
+
+Every online fsck scrubbing function is expected to read every ondisk metadata
+block of a structure in the course of checking the structure.
+Corruption problems observed during a check are immediately reported to
+userspace as corruption; during a cross-reference, they are reported as a
+failure to cross-reference once the full examination is complete.
+Reads satisfied by a buffer already in cache (and hence already verified)
+bypass these checks.
+
+Internal Consistency Checks
+```````````````````````````
+
+After the buffer cache, the next level of metadata protection is the internal
+record verification code built into the filesystem.
+These checks are split between the buffer verifiers, the in-filesystem users of
+the buffer cache, and the scrub code itself, depending on the amount of higher
+level context required.
+The scope of checking is still internal to the block.
+These higher level checking functions answer these questions:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- If the block contains records, do the records fit within the block?
+
+- If the block tracks internal free space information, is it consistent with
+  the record areas?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+Record checks in this category are more rigorous and more time-intensive.
+For example, block pointers and inumbers are checked to ensure that they point
+within the dynamically allocated parts of an allocation group and within
+the filesystem.
+Names are checked for invalid characters, and flags are checked for invalid
+combinations.
+Other record attributes are checked for sensible values.
+Btree records spanning an interval of the btree keyspace are checked for
+correct order and lack of mergeability (except for file fork mappings).
+For performance reasons, regular code may skip some of these checks unless
+debugging is enabled or a write is about to occur.
+Scrub functions, of course, must check all possible problems.
+
+Validation of Userspace-Controlled Record Attributes
+````````````````````````````````````````````````````
+
+Various pieces of filesystem metadata are directly controlled by userspace.
+Because of this nature, validation work cannot be more precise than checking
+that a value is within the possible range.
+These fields include:
+
+- Superblock fields controlled by mount options
+- Filesystem labels
+- File timestamps
+- File permissions
+- File size
+- File flags
+- Names present in directory entries, extended attribute keys, and filesystem
+  labels
+- Extended attribute key namespaces
+- Extended attribute values
+- File data block contents
+- Quota limits
+- Quota timer expiration (if resource usage exceeds the soft limit)
+
+Cross-Referencing Space Metadata
+````````````````````````````````
+
+After internal block checks, the next higher level of checking is
+cross-referencing records between metadata structures.
+For regular runtime code, the cost of these checks is considered to be
+prohibitively expensive, but as scrub is dedicated to rooting out
+inconsistencies, it must pursue all avenues of inquiry.
+The exact set of cross-referencing is highly dependent on the context of the
+data structure being checked.
+
+The XFS btree code has keyspace scanning functions that online fsck uses to
+cross reference one structure with another.
+Specifically, scrub can scan the key space of an index to determine if that
+keyspace is fully, sparsely, or not at all mapped to records.
+For the reverse mapping btree, it is possible to mask parts of the key for the
+purposes of performing a keyspace scan so that scrub can decide if the rmap
+btree contains records mapping a certain extent of physical space without the
+sparsenses of the rest of the rmap keyspace getting in the way.
+
+Btree blocks undergo the following checks before cross-referencing:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- Do the records fit within the block?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+- Are the name hashes in the correct order?
+
+- Do node pointers within the btree point to valid block addresses for the type
+  of btree?
+
+- Do child pointers point towards the leaves?
+
+- Do sibling pointers point across the same level?
+
+- For each node block record, does the record key accurate reflect the contents
+  of the child block?
+
+Space allocation records are cross-referenced as follows:
+
+1. Any space mentioned by any metadata structure are cross-referenced as
+   follows:
+
+   - Does the reverse mapping index list only the appropriate owner as the
+     owner of each block?
+
+   - Are none of the blocks claimed as free space?
+
+   - If these aren't file data blocks, are none of the blocks claimed as space
+     shared by different owners?
+
+2. Btree blocks are cross-referenced as follows:
+
+   - Everything in class 1 above.
+
+   - If there's a parent node block, do the keys listed for this block match the
+     keyspace of this block?
+
+   - Do the sibling pointers point to valid blocks?  Of the same level?
+
+   - Do the child pointers point to valid blocks?  Of the next level down?
+
+3. Free space btree records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Does the reverse mapping index list no owners of this space?
+
+   - Is this space not claimed by the inode index for inodes?
+
+   - Is it not mentioned by the reference count index?
+
+   - Is there a matching record in the other free space btree?
+
+4. Inode btree records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Is there a matching record in free inode btree?
+
+   - Do cleared bits in the holemask correspond with inode clusters?
+
+   - Do set bits in the freemask correspond with inode records with zero link
+     count?
+
+5. Inode records are cross-referenced as follows:
+
+   - Everything in class 1.
+
+   - Do all the fields that summarize information about the file forks actually
+     match those forks?
+
+   - Does each inode with zero link count correspond to a record in the free
+     inode btree?
+
+6. File fork space mapping records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Is this space not mentioned by the inode btrees?
+
+   - If this is a CoW fork mapping, does it correspond to a CoW entry in the
+     reference count btree?
+
+7. Reference count records are cross-referenced as follows:
+
+   - Everything in class 1 and 2 above.
+
+   - Within the space subkeyspace of the rmap btree (that is to say, all
+     records mapped to a particular space extent and ignoring the owner info),
+     are there the same number of reverse mapping records for each block as the
+     reference count record claims?
+
+Checking Extended Attributes
+````````````````````````````
+
+Extended attributes implement a key-value store that enable fragments of data
+to be attached to any file.
+Both the kernel and userspace can access the keys and values, subject to
+namespace and privilege restrictions.
+Most typically these fragments are metadata about the file -- origins, security
+contexts, user-supplied labels, indexing information, etc.
+
+Names can be as long as 255 bytes and can exist in several different
+namespaces.
+Values can be as large as 64KB.
+A file's extended attributes are stored in blocks mapped by the attr fork.
+The mappings point to leaf blocks, remote value blocks, or dabtree blocks.
+Block 0 in the attribute fork is always the top of the structure, but otherwise
+each of the three types of blocks can be found at any offset in the attr fork.
+Leaf blocks contain attribute key records that point to the name and the value.
+Names are always stored elsewhere in the same leaf block.
+Values that are less than 3/4 the size of a filesystem block are also stored
+elsewhere in the same leaf block.
+Remote value blocks contain values that are too large to fit inside a leaf.
+If the leaf information exceeds a single filesystem block, a dabtree (also
+rooted at block 0) is created to map hashes of the attribute names to leaf
+blocks in the attr fork.
+
+Checking an extended attribute structure is not so straightforward due to the
+lack of separation between attr blocks and index blocks.
+Scrub must read each block mapped by the attr fork and ignore the non-leaf
+blocks:
+
+1. Walk the dabtree in the attr fork (if present) to ensure that there are no
+   irregularities in the blocks or dabtree mappings that do not point to
+   attr leaf blocks.
+
+2. Walk the blocks of the attr fork looking for leaf blocks.
+   For each entry inside a leaf:
+
+   a. Validate that the name does not contain invalid characters.
+
+   b. Read the attr value.
+      This performs a named lookup of the attr name to ensure the correctness
+      of the dabtree.
+      If the value is stored in a remote block, this also validates the
+      integrity of the remote value block.
+
+Checking and Cross-Referencing Directories
+``````````````````````````````````````````
+
+The filesystem directory tree is a directed acylic graph structure, with files
+constituting the nodes, and directory entries (dirents) constituting the edges.
+Directories are a special type of file containing a set of mappings from a
+255-byte sequence (name) to an inumber.
+These are called directory entries, or dirents for short.
+Each directory file must have exactly one directory pointing to the file.
+A root directory points to itself.
+Directory entries point to files of any type.
+Each non-directory file may have multiple directories point to it.
+
+In XFS, directories are implemented as a file containing up to three 32GB
+partitions.
+The first partition contains directory entry data blocks.
+Each data block contains variable-sized records associating a user-provided
+name with an inumber and, optionally, a file type.
+If the directory entry data grows beyond one block, the second partition (which
+exists as post-EOF extents) is populated with a block containing free space
+information and an index that maps hashes of the dirent names to directory data
+blocks in the first partition.
+This makes directory name lookups very fast.
+If this second partition grows beyond one block, the third partition is
+populated with a linear array of free space information for faster
+expansions.
+If the free space has been separated and the second partition grows again
+beyond one block, then a dabtree is used to map hashes of dirent names to
+directory data blocks.
+
+Checking a directory is pretty straightforward:
+
+1. Walk the dabtree in the second partition (if present) to ensure that there
+   are no irregularities in the blocks or dabtree mappings that do not point to
+   dirent blocks.
+
+2. Walk the blocks of the first partition looking for directory entries.
+   Each dirent is checked as follows:
+
+   a. Does the name contain no invalid characters?
+
+   b. Does the inumber correspond to an actual, allocated inode?
+
+   c. Does the child inode have a nonzero link count?
+
+   d. If a file type is included in the dirent, does it match the type of the
+      inode?
+
+   e. If the child is a subdirectory, does the child's dotdot pointer point
+      back to the parent?
+
+   f. If the directory has a second partition, perform a named lookup of the
+      dirent name to ensure the correctness of the dabtree.
+
+3. Walk the free space list in the third partition (if present) to ensure that
+   the free spaces it describes are really unused.
+
+Checking operations involving :ref:`parents <dirparent>` and
+:ref:`file link counts <nlinks>` are discussed in more detail in later
+sections.
+
+Checking Directory/Attribute Btrees
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+As stated in previous sections, the directory/attribute btree (dabtree) index
+maps user-provided names to improve lookup times by avoiding linear scans.
+Internally, it maps a 32-bit hash of the name to a block offset within the
+appropriate file fork.
+
+The internal structure of a dabtree closely resembles the btrees that record
+fixed-size metadata records -- each dabtree block contains a magic number, a
+checksum, sibling pointers, a UUID, a tree level, and a log sequence number.
+The format of leaf and node records are the same -- each entry points to the
+next level down in the hierarchy, with dabtree node records pointing to dabtree
+leaf blocks, and dabtree leaf records pointing to non-dabtree blocks elsewhere
+in the fork.
+
+Checking and cross-referencing the dabtree is very similar to what is done for
+space btrees:
+
+- Does the type of data stored in the block match what scrub is expecting?
+
+- Does the block belong to the owning structure that asked for the read?
+
+- Do the records fit within the block?
+
+- Are the records contained inside the block free of obvious corruptions?
+
+- Are the name hashes in the correct order?
+
+- Do node pointers within the dabtree point to valid fork offsets for dabtree
+  blocks?
+
+- Do leaf pointers within the dabtree point to valid fork offsets for directory
+  or attr leaf blocks?
+
+- Do child pointers point towards the leaves?
+
+- Do sibling pointers point across the same level?
+
+- For each dabtree node record, does the record key accurate reflect the
+  contents of the child dabtree block?
+
+- For each dabtree leaf record, does the record key accurate reflect the
+  contents of the directory or attr block?
+
+Cross-Referencing Summary Counters
+``````````````````````````````````
+
+XFS maintains three classes of summary counters: available resources, quota
+resource usage, and file link counts.
+
+In theory, the amount of available resources (data blocks, inodes, realtime
+extents) can be found by walking the entire filesystem.
+This would make for very slow reporting, so a transactional filesystem can
+maintain summaries of this information in the superblock.
+Cross-referencing these values against the filesystem metadata should be a
+simple matter of walking the free space and inode metadata in each AG and the
+realtime bitmap, but there are complications that will be discussed in
+:ref:`more detail <fscounters>` later.
+
+:ref:`Quota usage <quotacheck>` and :ref:`file link count <nlinks>`
+checking are sufficiently complicated to warrant separate sections.
+
+Post-Repair Reverification
+``````````````````````````
+
+After performing a repair, the checking code is run a second time to validate
+the new structure, and the results of the health assessment are recorded
+internally and returned to the calling process.
+This step is critical for enabling system administrator to monitor the status
+of the filesystem and the progress of any repairs.
+For developers, it is a useful means to judge the efficacy of error detection
+and correction in the online and offline checking tools.
+
+Eventual Consistency vs. Online Fsck
+------------------------------------
+
+Complex operations can make modifications to multiple per-AG data structures
+with a chain of transactions.
+These chains, once committed to the log, are restarted during log recovery if
+the system crashes while processing the chain.
+Because the AG header buffers are unlocked between transactions within a chain,
+online checking must coordinate with chained operations that are in progress to
+avoid incorrectly detecting inconsistencies due to pending chains.
+Furthermore, online repair must not run when operations are pending because
+the metadata are temporarily inconsistent with each other, and rebuilding is
+not possible.
+
+Only online fsck has this requirement of total consistency of AG metadata, and
+should be relatively rare as compared to filesystem change operations.
+Online fsck coordinates with transaction chains as follows:
+
+* For each AG, maintain a count of intent items targeting that AG.
+  The count should be bumped whenever a new item is added to the chain.
+  The count should be dropped when the filesystem has locked the AG header
+  buffers and finished the work.
+
+* When online fsck wants to examine an AG, it should lock the AG header
+  buffers to quiesce all transaction chains that want to modify that AG.
+  If the count is zero, proceed with the checking operation.
+  If it is nonzero, cycle the buffer locks to allow the chain to make forward
+  progress.
+
+This may lead to online fsck taking a long time to complete, but regular
+filesystem updates take precedence over background checking activity.
+Details about the discovery of this situation are presented in the
+:ref:`next section <chain_coordination>`, and details about the solution
+are presented :ref:`after that<intent_drains>`.
+
+.. _chain_coordination:
+
+Discovery of the Problem
+````````````````````````
+
+Midway through the development of online scrubbing, the fsstress tests
+uncovered a misinteraction between online fsck and compound transaction chains
+created by other writer threads that resulted in false reports of metadata
+inconsistency.
+The root cause of these reports is the eventual consistency model introduced by
+the expansion of deferred work items and compound transaction chains when
+reverse mapping and reflink were introduced.
+
+Originally, transaction chains were added to XFS to avoid deadlocks when
+unmapping space from files.
+Deadlock avoidance rules require that AGs only be locked in increasing order,
+which makes it impossible (say) to use a single transaction to free a space
+extent in AG 7 and then try to free a now superfluous block mapping btree block
+in AG 3.
+To avoid these kinds of deadlocks, XFS creates Extent Freeing Intent (EFI) log
+items to commit to freeing some space in one transaction while deferring the
+actual metadata updates to a fresh transaction.
+The transaction sequence looks like this:
+
+1. The first transaction contains a physical update to the file's block mapping
+   structures to remove the mapping from the btree blocks.
+   It then attaches to the in-memory transaction an action item to schedule
+   deferred freeing of space.
+   Concretely, each transaction maintains a list of ``struct
+   xfs_defer_pending`` objects, each of which maintains a list of ``struct
+   xfs_extent_free_item`` objects.
+   Returning to the example above, the action item tracks the freeing of both
+   the unmapped space from AG 7 and the block mapping btree (BMBT) block from
+   AG 3.
+   Deferred frees recorded in this manner are committed in the log by creating
+   an EFI log item from the ``struct xfs_extent_free_item`` object and
+   attaching the log item to the transaction.
+   When the log is persisted to disk, the EFI item is written into the ondisk
+   transaction record.
+   EFIs can list up to 16 extents to free, all sorted in AG order.
+
+2. The second transaction contains a physical update to the free space btrees
+   of AG 3 to release the former BMBT block and a second physical update to the
+   free space btrees of AG 7 to release the unmapped file space.
+   Observe that the physical updates are resequenced in the correct order
+   when possible.
+   Attached to the transaction is a an extent free done (EFD) log item.
+   The EFD contains a pointer to the EFI logged in transaction #1 so that log
+   recovery can tell if the EFI needs to be replayed.
+
+If the system goes down after transaction #1 is written back to the filesystem
+but before #2 is committed, a scan of the filesystem metadata would show
+inconsistent filesystem metadata because there would not appear to be any owner
+of the unmapped space.
+Happily, log recovery corrects this inconsistency for us -- when recovery finds
+an intent log item but does not find a corresponding intent done item, it will
+reconstruct the incore state of the intent item and finish it.
+In the example above, the log must replay both frees described in the recovered
+EFI to complete the recovery phase.
+
+There are subtleties to XFS' transaction chaining strategy to consider:
+
+* Log items must be added to a transaction in the correct order to prevent
+  conflicts with principal objects that are not held by the transaction.
+  In other words, all per-AG metadata updates for an unmapped block must be
+  completed before the last update to free the extent, and extents should not
+  be reallocated until that last update commits to the log.
+
+* AG header buffers are released between each transaction in a chain.
+  This means that other threads can observe an AG in an intermediate state,
+  but as long as the first subtlety is handled, this should not affect the
+  correctness of filesystem operations.
+
+* Unmounting the filesystem flushes all pending work to disk, which means that
+  offline fsck never sees the temporary inconsistencies caused by deferred
+  work item processing.
+
+In this manner, XFS employs a form of eventual consistency to avoid deadlocks
+and increase parallelism.
+
+During the design phase of the reverse mapping and reflink features, it was
+decided that it was impractical to cram all the reverse mapping updates for a
+single filesystem change into a single transaction because a single file
+mapping operation can explode into many small updates:
+
+* The block mapping update itself
+* A reverse mapping update for the block mapping update
+* Fixing the freelist
+* A reverse mapping update for the freelist fix
+
+* A shape change to the block mapping btree
+* A reverse mapping update for the btree update
+* Fixing the freelist (again)
+* A reverse mapping update for the freelist fix
+
+* An update to the reference counting information
+* A reverse mapping update for the refcount update
+* Fixing the freelist (a third time)
+* A reverse mapping update for the freelist fix
+
+* Freeing any space that was unmapped and not owned by any other file
+* Fixing the freelist (a fourth time)
+* A reverse mapping update for the freelist fix
+
+* Freeing the space used by the block mapping btree
+* Fixing the freelist (a fifth time)
+* A reverse mapping update for the freelist fix
+
+Free list fixups are not usually needed more than once per AG per transaction
+chain, but it is theoretically possible if space is very tight.
+For copy-on-write updates this is even worse, because this must be done once to
+remove the space from a staging area and again to map it into the file!
+
+To deal with this explosion in a calm manner, XFS expands its use of deferred
+work items to cover most reverse mapping updates and all refcount updates.
+This reduces the worst case size of transaction reservations by breaking the
+work into a long chain of small updates, which increases the degree of eventual
+consistency in the system.
+Again, this generally isn't a problem because XFS orders its deferred work
+items carefully to avoid resource reuse conflicts between unsuspecting threads.
+
+However, online fsck changes the rules -- remember that although physical
+updates to per-AG structures are coordinated by locking the buffers for AG
+headers, buffer locks are dropped between transactions.
+Once scrub acquires resources and takes locks for a data structure, it must do
+all the validation work without releasing the lock.
+If the main lock for a space btree is an AG header buffer lock, scrub may have
+interrupted another thread that is midway through finishing a chain.
+For example, if a thread performing a copy-on-write has completed a reverse
+mapping update but not the corresponding refcount update, the two AG btrees
+will appear inconsistent to scrub and an observation of corruption will be
+recorded.  This observation will not be correct.
+If a repair is attempted in this state, the results will be catastrophic!
+
+Several other solutions to this problem were evaluated upon discovery of this
+flaw and rejected:
+
+1. Add a higher level lock to allocation groups and require writer threads to
+   acquire the higher level lock in AG order before making any changes.
+   This would be very difficult to implement in practice because it is
+   difficult to determine which locks need to be obtained, and in what order,
+   without simulating the entire operation.
+   Performing a dry run of a file operation to discover necessary locks would
+   make the filesystem very slow.
+
+2. Make the deferred work coordinator code aware of consecutive intent items
+   targeting the same AG and have it hold the AG header buffers locked across
+   the transaction roll between updates.
+   This would introduce a lot of complexity into the coordinator since it is
+   only loosely coupled with the actual deferred work items.
+   It would also fail to solve the problem because deferred work items can
+   generate new deferred subtasks, but all subtasks must be complete before
+   work can start on a new sibling task.
+
+3. Teach online fsck to walk all transactions waiting for whichever lock(s)
+   protect the data structure being scrubbed to look for pending operations.
+   The checking and repair operations must factor these pending operations into
+   the evaluations being performed.
+   This solution is a nonstarter because it is *extremely* invasive to the main
+   filesystem.
+
+.. _intent_drains:
+
+Intent Drains
+`````````````
+
+Online fsck uses an atomic intent item counter and lock cycling to coordinate
+with transaction chains.
+There are two key properties to the drain mechanism.
+First, the counter is incremented when a deferred work item is *queued* to a
+transaction, and it is decremented after the associated intent done log item is
+*committed* to another transaction.
+The second property is that deferred work can be added to a transaction without
+holding an AG header lock, but per-AG work items cannot be marked done without
+locking that AG header buffer to log the physical updates and the intent done
+log item.
+The first property enables scrub to yield to running transaction chains, which
+is an explicit deprioritization of online fsck to benefit file operations.
+The second property of the drain is key to the correct coordination of scrub,
+since scrub will always be able to decide if a conflict is possible.
+
+For regular filesystem code, the drain works as follows:
+
+1. Call the appropriate subsystem function to add a deferred work item to a
+   transaction.
+
+2. The function calls ``xfs_defer_drain_bump`` to increase the counter.
+
+3. When the deferred item manager wants to finish the deferred work item, it
+   calls ``->finish_item`` to complete it.
+
+4. The ``->finish_item`` implementation logs some changes and calls
+   ``xfs_defer_drain_drop`` to decrease the sloppy counter and wake up any threads
+   waiting on the drain.
+
+5. The subtransaction commits, which unlocks the resource associated with the
+   intent item.
+
+For scrub, the drain works as follows:
+
+1. Lock the resource(s) associated with the metadata being scrubbed.
+   For example, a scan of the refcount btree would lock the AGI and AGF header
+   buffers.
+
+2. If the counter is zero (``xfs_defer_drain_busy`` returns false), there are no
+   chains in progress and the operation may proceed.
+
+3. Otherwise, release the resources grabbed in step 1.
+
+4. Wait for the intent counter to reach zero (``xfs_defer_drain_intents``), then go
+   back to step 1 unless a signal has been caught.
+
+To avoid polling in step 4, the drain provides a waitqueue for scrub threads to
+be woken up whenever the intent count drops to zero.
+
+.. _jump_labels:
+
+Static Keys (aka Jump Label Patching)
+`````````````````````````````````````
+
+Online fsck for XFS separates the regular filesystem from the checking and
+repair code as much as possible.
+However, there are a few parts of online fsck (such as the intent drains, and
+later, live update hooks) where it is useful for the online fsck code to know
+what's going on in the rest of the filesystem.
+Since it is not expected that online fsck will be constantly running in the
+background, it is very important to minimize the runtime overhead imposed by
+these hooks when online fsck is compiled into the kernel but not actively
+running on behalf of userspace.
+Taking locks in the hot path of a writer thread to access a data structure only
+to find that no further action is necessary is expensive -- on the author's
+computer, this have an overhead of 40-50ns per access.
+Fortunately, the kernel supports dynamic code patching, which enables XFS to
+replace a static branch to hook code with ``nop`` sleds when online fsck isn't
+running.
+This sled has an overhead of however long it takes the instruction decoder to
+skip past the sled, which seems to be on the order of less than 1ns and
+does not access memory outside of instruction fetching.
+
+When online fsck enables the static key, the sled is replaced with an
+unconditional branch to call the hook code.
+The switchover is quite expensive (~22000ns) but is paid entirely by the
+program that invoked online fsck, and can be amortized if multiple threads
+enter online fsck at the same time, or if multiple filesystems are being
+checked at the same time.
+Changing the branch direction requires taking the CPU hotplug lock, and since
+CPU initialization requires memory allocation, online fsck must be careful not
+to change a static key while holding any locks or resources that could be
+accessed in the memory reclaim paths.
+To minimize contention on the CPU hotplug lock, care should be taken not to
+enable or disable static keys unnecessarily.
+
+Because static keys are intended to minimize hook overhead for regular
+filesystem operations when xfs_scrub is not running, the intended usage
+patterns are as follows:
+
+- The hooked part of XFS should declare a static-scoped static key that
+  defaults to false.
+  The ``DEFINE_STATIC_KEY_FALSE`` macro takes care of this.
+  The static key itself should be declared as a ``static`` variable.
+
+- When deciding to invoke code that's only used by scrub, the regular
+  filesystem should call the ``static_branch_unlikely`` predicate to avoid the
+  scrub-only hook code if the static key is not enabled.
+
+- The regular filesystem should export helper functions that call
+  ``static_branch_inc`` to enable and ``static_branch_dec`` to disable the
+  static key.
+  Wrapper functions make it easy to compile out the relevant code if the kernel
+  distributor turns off online fsck at build time.
+
+- Scrub functions wanting to turn on scrub-only XFS functionality should call
+  the ``xchk_fsgates_enable`` from the setup function to enable a specific
+  hook.
+  This must be done before obtaining any resources that are used by memory
+  reclaim.
+  Callers had better be sure they really need the functionality gated by the
+  static key; the ``TRY_HARDER`` flag is useful here.
+
+Online scrub has resource acquisition helpers (e.g. ``xchk_perag_lock``) to
+handle locking AGI and AGF buffers for all scrubber functions.
+If it detects a conflict between scrub and the running transactions, it will
+try to wait for intents to complete.
+If the caller of the helper has not enabled the static key, the helper will
+return -EDEADLOCK, which should result in the scrub being restarted with the
+``TRY_HARDER`` flag set.
+The scrub setup function should detect that flag, enable the static key, and
+try the scrub again.
+Scrub teardown disables all static keys obtained by ``xchk_fsgates_enable``.
+
+For more information, please see the kernel documentation of
+Documentation/staging/static-keys.rst.
+
+.. _xfile:
+
+Pageable Kernel Memory
+----------------------
+
+Some online checking functions work by scanning the filesystem to build a
+shadow copy of an ondisk metadata structure in memory and comparing the two
+copies.
+For online repair to rebuild a metadata structure, it must compute the record
+set that will be stored in the new structure before it can persist that new
+structure to disk.
+Ideally, repairs complete with a single atomic commit that introduces
+a new data structure.
+To meet these goals, the kernel needs to collect a large amount of information
+in a place that doesn't require the correct operation of the filesystem.
+
+Kernel memory isn't suitable because:
+
+* Allocating a contiguous region of memory to create a C array is very
+  difficult, especially on 32-bit systems.
+
+* Linked lists of records introduce double pointer overhead which is very high
+  and eliminate the possibility of indexed lookups.
+
+* Kernel memory is pinned, which can drive the system into OOM conditions.
+
+* The system might not have sufficient memory to stage all the information.
+
+At any given time, online fsck does not need to keep the entire record set in
+memory, which means that individual records can be paged out if necessary.
+Continued development of online fsck demonstrated that the ability to perform
+indexed data storage would also be very useful.
+Fortunately, the Linux kernel already has a facility for byte-addressable and
+pageable storage: tmpfs.
+In-kernel graphics drivers (most notably i915) take advantage of tmpfs files
+to store intermediate data that doesn't need to be in memory at all times, so
+that usage precedent is already established.
+Hence, the ``xfile`` was born!
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| The first edition of online repair inserted records into a new btree as  |
+| it found them, which failed because filesystem could shut down with a    |
+| built data structure, which would be live after recovery finished.       |
+|                                                                          |
+| The second edition solved the half-rebuilt structure problem by storing  |
+| everything in memory, but frequently ran the system out of memory.       |
+|                                                                          |
+| The third edition solved the OOM problem by using linked lists, but the  |
+| memory overhead of the list pointers was extreme.                        |
++--------------------------------------------------------------------------+
+
+xfile Access Models
+```````````````````
+
+A survey of the intended uses of xfiles suggested these use cases:
+
+1. Arrays of fixed-sized records (space management btrees, directory and
+   extended attribute entries)
+
+2. Sparse arrays of fixed-sized records (quotas and link counts)
+
+3. Large binary objects (BLOBs) of variable sizes (directory and extended
+   attribute names and values)
+
+4. Staging btrees in memory (reverse mapping btrees)
+
+5. Arbitrary contents (realtime space management)
+
+To support the first four use cases, high level data structures wrap the xfile
+to share functionality between online fsck functions.
+The rest of this section discusses the interfaces that the xfile presents to
+four of those five higher level data structures.
+The fifth use case is discussed in the :ref:`realtime summary <rtsummary>` case
+study.
+
+XFS is very record-based, which suggests that the ability to load and store
+complete records is important.
+To support these cases, a pair of ``xfile_load`` and ``xfile_store``
+functions are provided to read and persist objects into an xfile that treat any
+error as an out of memory error.  For online repair, squashing error conditions
+in this manner is an acceptable behavior because the only reaction is to abort
+the operation back to userspace.
+
+However, no discussion of file access idioms is complete without answering the
+question, "But what about mmap?"
+It is convenient to access storage directly with pointers, just like userspace
+code does with regular memory.
+Online fsck must not drive the system into OOM conditions, which means that
+xfiles must be responsive to memory reclamation.
+tmpfs can only push a pagecache folio to the swap cache if the folio is neither
+pinned nor locked, which means the xfile must not pin too many folios.
+
+Short term direct access to xfile contents is done by locking the pagecache
+folio and mapping it into kernel address space.  Object load and store uses this
+mechanism.  Folio locks are not supposed to be held for long periods of time, so
+long term direct access to xfile contents is done by bumping the folio refcount,
+mapping it into kernel address space, and dropping the folio lock.
+These long term users *must* be responsive to memory reclaim by hooking into
+the shrinker infrastructure to know when to release folios.
+
+The ``xfile_get_folio`` and ``xfile_put_folio`` functions are provided to
+retrieve the (locked) folio that backs part of an xfile and to release it.
+The only code to use these folio lease functions are the xfarray
+:ref:`sorting<xfarray_sort>` algorithms and the :ref:`in-memory
+btrees<xfbtree>`.
+
+xfile Access Coordination
+`````````````````````````
+
+For security reasons, xfiles must be owned privately by the kernel.
+They are marked ``S_PRIVATE`` to prevent interference from the security system,
+must never be mapped into process file descriptor tables, and their pages must
+never be mapped into userspace processes.
+
+To avoid locking recursion issues with the VFS, all accesses to the shmfs file
+are performed by manipulating the page cache directly.
+xfile writers call the ``->write_begin`` and ``->write_end`` functions of the
+xfile's address space to grab writable pages, copy the caller's buffer into the
+page, and release the pages.
+xfile readers call ``shmem_read_mapping_page_gfp`` to grab pages directly
+before copying the contents into the caller's buffer.
+In other words, xfiles ignore the VFS read and write code paths to avoid
+having to create a dummy ``struct kiocb`` and to avoid taking inode and
+freeze locks.
+tmpfs cannot be frozen, and xfiles must not be exposed to userspace.
+
+If an xfile is shared between threads to stage repairs, the caller must provide
+its own locks to coordinate access.
+For example, if a scrub function stores scan results in an xfile and needs
+other threads to provide updates to the scanned data, the scrub function must
+provide a lock for all threads to share.
+
+.. _xfarray:
+
+Arrays of Fixed-Sized Records
+`````````````````````````````
+
+In XFS, each type of indexed space metadata (free space, inodes, reference
+counts, file fork space, and reverse mappings) consists of a set of fixed-size
+records indexed with a classic B+ tree.
+Directories have a set of fixed-size dirent records that point to the names,
+and extended attributes have a set of fixed-size attribute keys that point to
+names and values.
+Quota counters and file link counters index records with numbers.
+During a repair, scrub needs to stage new records during the gathering step and
+retrieve them during the btree building step.
+
+Although this requirement can be satisfied by calling the read and write
+methods of the xfile directly, it is simpler for callers for there to be a
+higher level abstraction to take care of computing array offsets, to provide
+iterator functions, and to deal with sparse records and sorting.
+The ``xfarray`` abstraction presents a linear array for fixed-size records atop
+the byte-accessible xfile.
+
+.. _xfarray_access_patterns:
+
+Array Access Patterns
+^^^^^^^^^^^^^^^^^^^^^
+
+Array access patterns in online fsck tend to fall into three categories.
+Iteration of records is assumed to be necessary for all cases and will be
+covered in the next section.
+
+The first type of caller handles records that are indexed by position.
+Gaps may exist between records, and a record may be updated multiple times
+during the collection step.
+In other words, these callers want a sparse linearly addressed table file.
+The typical use case are quota records or file link count records.
+Access to array elements is performed programmatically via ``xfarray_load`` and
+``xfarray_store`` functions, which wrap the similarly-named xfile functions to
+provide loading and storing of array elements at arbitrary array indices.
+Gaps are defined to be null records, and null records are defined to be a
+sequence of all zero bytes.
+Null records are detected by calling ``xfarray_element_is_null``.
+They are created either by calling ``xfarray_unset`` to null out an existing
+record or by never storing anything to an array index.
+
+The second type of caller handles records that are not indexed by position
+and do not require multiple updates to a record.
+The typical use case here is rebuilding space btrees and key/value btrees.
+These callers can add records to the array without caring about array indices
+via the ``xfarray_append`` function, which stores a record at the end of the
+array.
+For callers that require records to be presentable in a specific order (e.g.
+rebuilding btree data), the ``xfarray_sort`` function can arrange the sorted
+records; this function will be covered later.
+
+The third type of caller is a bag, which is useful for counting records.
+The typical use case here is constructing space extent reference counts from
+reverse mapping information.
+Records can be put in the bag in any order, they can be removed from the bag
+at any time, and uniqueness of records is left to callers.
+The ``xfarray_store_anywhere`` function is used to insert a record in any
+null record slot in the bag; and the ``xfarray_unset`` function removes a
+record from the bag.
+
+Iterating Array Elements
+^^^^^^^^^^^^^^^^^^^^^^^^
+
+Most users of the xfarray require the ability to iterate the records stored in
+the array.
+Callers can probe every possible array index with the following:
+
+.. code-block:: c
+
+	xfarray_idx_t i;
+	foreach_xfarray_idx(array, i) {
+	    xfarray_load(array, i, &rec);
+
+	    /* do something with rec */
+	}
+
+All users of this idiom must be prepared to handle null records or must already
+know that there aren't any.
+
+For xfarray users that want to iterate a sparse array, the ``xfarray_iter``
+function ignores indices in the xfarray that have never been written to by
+calling ``xfile_seek_data`` (which internally uses ``SEEK_DATA``) to skip areas
+of the array that are not populated with memory pages.
+Once it finds a page, it will skip the zeroed areas of the page.
+
+.. code-block:: c
+
+	xfarray_idx_t i = XFARRAY_CURSOR_INIT;
+	while ((ret = xfarray_iter(array, &i, &rec)) == 1) {
+	    /* do something with rec */
+	}
+
+.. _xfarray_sort:
+
+Sorting Array Elements
+^^^^^^^^^^^^^^^^^^^^^^
+
+During the fourth demonstration of online repair, a community reviewer remarked
+that for performance reasons, online repair ought to load batches of records
+into btree record blocks instead of inserting records into a new btree one at a
+time.
+The btree insertion code in XFS is responsible for maintaining correct ordering
+of the records, so naturally the xfarray must also support sorting the record
+set prior to bulk loading.
+
+Case Study: Sorting xfarrays
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+The sorting algorithm used in the xfarray is actually a combination of adaptive
+quicksort and a heapsort subalgorithm in the spirit of
+`Sedgewick <https://algs4.cs.princeton.edu/23quicksort/>`_ and
+`pdqsort <https://github.com/orlp/pdqsort>`_, with customizations for the Linux
+kernel.
+To sort records in a reasonably short amount of time, ``xfarray`` takes
+advantage of the binary subpartitioning offered by quicksort, but it also uses
+heapsort to hedge against performance collapse if the chosen quicksort pivots
+are poor.
+Both algorithms are (in general) O(n * lg(n)), but there is a wide performance
+gulf between the two implementations.
+
+The Linux kernel already contains a reasonably fast implementation of heapsort.
+It only operates on regular C arrays, which limits the scope of its usefulness.
+There are two key places where the xfarray uses it:
+
+* Sorting any record subset backed by a single xfile page.
+
+* Loading a small number of xfarray records from potentially disparate parts
+  of the xfarray into a memory buffer, and sorting the buffer.
+
+In other words, ``xfarray`` uses heapsort to constrain the nested recursion of
+quicksort, thereby mitigating quicksort's worst runtime behavior.
+
+Choosing a quicksort pivot is a tricky business.
+A good pivot splits the set to sort in half, leading to the divide and conquer
+behavior that is crucial to  O(n * lg(n)) performance.
+A poor pivot barely splits the subset at all, leading to O(n\ :sup:`2`)
+runtime.
+The xfarray sort routine tries to avoid picking a bad pivot by sampling nine
+records into a memory buffer and using the kernel heapsort to identify the
+median of the nine.
+
+Most modern quicksort implementations employ Tukey's "ninther" to select a
+pivot from a classic C array.
+Typical ninther implementations pick three unique triads of records, sort each
+of the triads, and then sort the middle value of each triad to determine the
+ninther value.
+As stated previously, however, xfile accesses are not entirely cheap.
+It turned out to be much more performant to read the nine elements into a
+memory buffer, run the kernel's in-memory heapsort on the buffer, and choose
+the 4th element of that buffer as the pivot.
+Tukey's ninthers are described in J. W. Tukey, `The ninther, a technique for
+low-effort robust (resistant) location in large samples`, in *Contributions to
+Survey Sampling and Applied Statistics*, edited by H. David, (Academic Press,
+1978), pp. 251–257.
+
+The partitioning of quicksort is fairly textbook -- rearrange the record
+subset around the pivot, then set up the current and next stack frames to
+sort with the larger and the smaller halves of the pivot, respectively.
+This keeps the stack space requirements to log2(record count).
+
+As a final performance optimization, the hi and lo scanning phase of quicksort
+keeps examined xfile pages mapped in the kernel for as long as possible to
+reduce map/unmap cycles.
+Surprisingly, this reduces overall sort runtime by nearly half again after
+accounting for the application of heapsort directly onto xfile pages.
+
+.. _xfblob:
+
+Blob Storage
+````````````
+
+Extended attributes and directories add an additional requirement for staging
+records: arbitrary byte sequences of finite length.
+Each directory entry record needs to store entry name,
+and each extended attribute needs to store both the attribute name and value.
+The names, keys, and values can consume a large amount of memory, so the
+``xfblob`` abstraction was created to simplify management of these blobs
+atop an xfile.
+
+Blob arrays provide ``xfblob_load`` and ``xfblob_store`` functions to retrieve
+and persist objects.
+The store function returns a magic cookie for every object that it persists.
+Later, callers provide this cookie to the ``xblob_load`` to recall the object.
+The ``xfblob_free`` function frees a specific blob, and the ``xfblob_truncate``
+function frees them all because compaction is not needed.
+
+The details of repairing directories and extended attributes will be discussed
+in a subsequent section about atomic file content exchanges.
+However, it should be noted that these repair functions only use blob storage
+to cache a small number of entries before adding them to a temporary ondisk
+file, which is why compaction is not required.
+
+.. _xfbtree:
+
+In-Memory B+Trees
+`````````````````
+
+The chapter about :ref:`secondary metadata<secondary_metadata>` mentioned that
+checking and repairing of secondary metadata commonly requires coordination
+between a live metadata scan of the filesystem and writer threads that are
+updating that metadata.
+Keeping the scan data up to date requires the ability to propagate
+metadata updates from the filesystem into the data being collected by the scan.
+This *can* be done by appending concurrent updates into a separate log file and
+applying them before writing the new metadata to disk, but this leads to
+unbounded memory consumption if the rest of the system is very busy.
+Another option is to skip the side-log and commit live updates from the
+filesystem directly into the scan data, which trades more overhead for a lower
+maximum memory requirement.
+In both cases, the data structure holding the scan results must support indexed
+access to perform well.
+
+Given that indexed lookups of scan data is required for both strategies, online
+fsck employs the second strategy of committing live updates directly into
+scan data.
+Because xfarrays are not indexed and do not enforce record ordering, they
+are not suitable for this task.
+Conveniently, however, XFS has a library to create and maintain ordered reverse
+mapping records: the existing rmap btree code!
+If only there was a means to create one in memory.
+
+Recall that the :ref:`xfile <xfile>` abstraction represents memory pages as a
+regular file, which means that the kernel can create byte or block addressable
+virtual address spaces at will.
+The XFS buffer cache specializes in abstracting IO to block-oriented  address
+spaces, which means that adaptation of the buffer cache to interface with
+xfiles enables reuse of the entire btree library.
+Btrees built atop an xfile are collectively known as ``xfbtrees``.
+The next few sections describe how they actually work.
+
+Using xfiles as a Buffer Cache Target
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Two modifications are necessary to support xfiles as a buffer cache target.
+The first is to make it possible for the ``struct xfs_buftarg`` structure to
+host the ``struct xfs_buf`` rhashtable, because normally those are held by a
+per-AG structure.
+The second change is to modify the buffer ``ioapply`` function to "read" cached
+pages from the xfile and "write" cached pages back to the xfile.
+Multiple access to individual buffers is controlled by the ``xfs_buf`` lock,
+since the xfile does not provide any locking on its own.
+With this adaptation in place, users of the xfile-backed buffer cache use
+exactly the same APIs as users of the disk-backed buffer cache.
+The separation between xfile and buffer cache implies higher memory usage since
+they do not share pages, but this property could some day enable transactional
+updates to an in-memory btree.
+Today, however, it simply eliminates the need for new code.
+
+Space Management with an xfbtree
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Space management for an xfile is very simple -- each btree block is one memory
+page in size.
+These blocks use the same header format as an on-disk btree, but the in-memory
+block verifiers ignore the checksums, assuming that xfile memory is no more
+corruption-prone than regular DRAM.
+Reusing existing code here is more important than absolute memory efficiency.
+
+The very first block of an xfile backing an xfbtree contains a header block.
+The header describes the owner, height, and the block number of the root
+xfbtree block.
+
+To allocate a btree block, use ``xfile_seek_data`` to find a gap in the file.
+If there are no gaps, create one by extending the length of the xfile.
+Preallocate space for the block with ``xfile_prealloc``, and hand back the
+location.
+To free an xfbtree block, use ``xfile_discard`` (which internally uses
+``FALLOC_FL_PUNCH_HOLE``) to remove the memory page from the xfile.
+
+Populating an xfbtree
+^^^^^^^^^^^^^^^^^^^^^
+
+An online fsck function that wants to create an xfbtree should proceed as
+follows:
+
+1. Call ``xfile_create`` to create an xfile.
+
+2. Call ``xfs_alloc_memory_buftarg`` to create a buffer cache target structure
+   pointing to the xfile.
+
+3. Pass the buffer cache target, buffer ops, and other information to
+   ``xfbtree_init`` to initialize the passed in ``struct xfbtree`` and write an
+   initial root block to the xfile.
+   Each btree type should define a wrapper that passes necessary arguments to
+   the creation function.
+   For example, rmap btrees define ``xfs_rmapbt_mem_create`` to take care of
+   all the necessary details for callers.
+
+4. Pass the xfbtree object to the btree cursor creation function for the
+   btree type.
+   Following the example above, ``xfs_rmapbt_mem_cursor`` takes care of this
+   for callers.
+
+5. Pass the btree cursor to the regular btree functions to make queries against
+   and to update the in-memory btree.
+   For example, a btree cursor for an rmap xfbtree can be passed to the
+   ``xfs_rmap_*`` functions just like any other btree cursor.
+   See the :ref:`next section<xfbtree_commit>` for information on dealing with
+   xfbtree updates that are logged to a transaction.
+
+6. When finished, delete the btree cursor, destroy the xfbtree object, free the
+   buffer target, and the destroy the xfile to release all resources.
+
+.. _xfbtree_commit:
+
+Committing Logged xfbtree Buffers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Although it is a clever hack to reuse the rmap btree code to handle the staging
+structure, the ephemeral nature of the in-memory btree block storage presents
+some challenges of its own.
+The XFS transaction manager must not commit buffer log items for buffers backed
+by an xfile because the log format does not understand updates for devices
+other than the data device.
+An ephemeral xfbtree probably will not exist by the time the AIL checkpoints
+log transactions back into the filesystem, and certainly won't exist during
+log recovery.
+For these reasons, any code updating an xfbtree in transaction context must
+remove the buffer log items from the transaction and write the updates into the
+backing xfile before committing or cancelling the transaction.
+
+The ``xfbtree_trans_commit`` and ``xfbtree_trans_cancel`` functions implement
+this functionality as follows:
+
+1. Find each buffer log item whose buffer targets the xfile.
+
+2. Record the dirty/ordered status of the log item.
+
+3. Detach the log item from the buffer.
+
+4. Queue the buffer to a special delwri list.
+
+5. Clear the transaction dirty flag if the only dirty log items were the ones
+   that were detached in step 3.
+
+6. Submit the delwri list to commit the changes to the xfile, if the updates
+   are being committed.
+
+After removing xfile logged buffers from the transaction in this manner, the
+transaction can be committed or cancelled.
+
+Bulk Loading of Ondisk B+Trees
+------------------------------
+
+As mentioned previously, early iterations of online repair built new btree
+structures by creating a new btree and adding observations individually.
+Loading a btree one record at a time had a slight advantage of not requiring
+the incore records to be sorted prior to commit, but was very slow and leaked
+blocks if the system went down during a repair.
+Loading records one at a time also meant that repair could not control the
+loading factor of the blocks in the new btree.
+
+Fortunately, the venerable ``xfs_repair`` tool had a more efficient means for
+rebuilding a btree index from a collection of records -- bulk btree loading.
+This was implemented rather inefficiently code-wise, since ``xfs_repair``
+had separate copy-pasted implementations for each btree type.
+
+To prepare for online fsck, each of the four bulk loaders were studied, notes
+were taken, and the four were refactored into a single generic btree bulk
+loading mechanism.
+Those notes in turn have been refreshed and are presented below.
+
+Geometry Computation
+````````````````````
+
+The zeroth step of bulk loading is to assemble the entire record set that will
+be stored in the new btree, and sort the records.
+Next, call ``xfs_btree_bload_compute_geometry`` to compute the shape of the
+btree from the record set, the type of btree, and any load factor preferences.
+This information is required for resource reservation.
+
+First, the geometry computation computes the minimum and maximum records that
+will fit in a leaf block from the size of a btree block and the size of the
+block header.
+Roughly speaking, the maximum number of records is::
+
+        maxrecs = (block_size - header_size) / record_size
+
+The XFS design specifies that btree blocks should be merged when possible,
+which means the minimum number of records is half of maxrecs::
+
+        minrecs = maxrecs / 2
+
+The next variable to determine is the desired loading factor.
+This must be at least minrecs and no more than maxrecs.
+Choosing minrecs is undesirable because it wastes half the block.
+Choosing maxrecs is also undesirable because adding a single record to each
+newly rebuilt leaf block will cause a tree split, which causes a noticeable
+drop in performance immediately afterwards.
+The default loading factor was chosen to be 75% of maxrecs, which provides a
+reasonably compact structure without any immediate split penalties::
+
+        default_load_factor = (maxrecs + minrecs) / 2
+
+If space is tight, the loading factor will be set to maxrecs to try to avoid
+running out of space::
+
+        leaf_load_factor = enough space ? default_load_factor : maxrecs
+
+Load factor is computed for btree node blocks using the combined size of the
+btree key and pointer as the record size::
+
+        maxrecs = (block_size - header_size) / (key_size + ptr_size)
+        minrecs = maxrecs / 2
+        node_load_factor = enough space ? default_load_factor : maxrecs
+
+Once that's done, the number of leaf blocks required to store the record set
+can be computed as::
+
+        leaf_blocks = ceil(record_count / leaf_load_factor)
+
+The number of node blocks needed to point to the next level down in the tree
+is computed as::
+
+        n_blocks = (n == 0 ? leaf_blocks : node_blocks[n])
+        node_blocks[n + 1] = ceil(n_blocks / node_load_factor)
+
+The entire computation is performed recursively until the current level only
+needs one block.
+The resulting geometry is as follows:
+
+- For AG-rooted btrees, this level is the root level, so the height of the new
+  tree is ``level + 1`` and the space needed is the summation of the number of
+  blocks on each level.
+
+- For inode-rooted btrees where the records in the top level do not fit in the
+  inode fork area, the height is ``level + 2``, the space needed is the
+  summation of the number of blocks on each level, and the inode fork points to
+  the root block.
+
+- For inode-rooted btrees where the records in the top level can be stored in
+  the inode fork area, then the root block can be stored in the inode, the
+  height is ``level + 1``, and the space needed is one less than the summation
+  of the number of blocks on each level.
+  This only becomes relevant when non-bmap btrees gain the ability to root in
+  an inode, which is a future patchset and only included here for completeness.
+
+.. _newbt:
+
+Reserving New B+Tree Blocks
+```````````````````````````
+
+Once repair knows the number of blocks needed for the new btree, it allocates
+those blocks using the free space information.
+Each reserved extent is tracked separately by the btree builder state data.
+To improve crash resilience, the reservation code also logs an Extent Freeing
+Intent (EFI) item in the same transaction as each space allocation and attaches
+its in-memory ``struct xfs_extent_free_item`` object to the space reservation.
+If the system goes down, log recovery will use the unfinished EFIs to free the
+unused space, the free space, leaving the filesystem unchanged.
+
+Each time the btree builder claims a block for the btree from a reserved
+extent, it updates the in-memory reservation to reflect the claimed space.
+Block reservation tries to allocate as much contiguous space as possible to
+reduce the number of EFIs in play.
+
+While repair is writing these new btree blocks, the EFIs created for the space
+reservations pin the tail of the ondisk log.
+It's possible that other parts of the system will remain busy and push the head
+of the log towards the pinned tail.
+To avoid livelocking the filesystem, the EFIs must not pin the tail of the log
+for too long.
+To alleviate this problem, the dynamic relogging capability of the deferred ops
+mechanism is reused here to commit a transaction at the log head containing an
+EFD for the old EFI and new EFI at the head.
+This enables the log to release the old EFI to keep the log moving forwards.
+
+EFIs have a role to play during the commit and reaping phases; please see the
+next section and the section about :ref:`reaping<reaping>` for more details.
+
+Writing the New Tree
+````````````````````
+
+This part is pretty simple -- the btree builder (``xfs_btree_bulkload``) claims
+a block from the reserved list, writes the new btree block header, fills the
+rest of the block with records, and adds the new leaf block to a list of
+written blocks::
+
+  ┌────┐
+  │leaf│
+  │RRR │
+  └────┘
+
+Sibling pointers are set every time a new block is added to the level::
+
+  ┌────┐ ┌────┐ ┌────┐ ┌────┐
+  │leaf│→│leaf│→│leaf│→│leaf│
+  │RRR │←│RRR │←│RRR │←│RRR │
+  └────┘ └────┘ └────┘ └────┘
+
+When it finishes writing the record leaf blocks, it moves on to the node
+blocks
+To fill a node block, it walks each block in the next level down in the tree
+to compute the relevant keys and write them into the parent node::
+
+      ┌────┐       ┌────┐
+      │node│──────→│node│
+      │PP  │←──────│PP  │
+      └────┘       └────┘
+      ↙   ↘         ↙   ↘
+  ┌────┐ ┌────┐ ┌────┐ ┌────┐
+  │leaf│→│leaf│→│leaf│→│leaf│
+  │RRR │←│RRR │←│RRR │←│RRR │
+  └────┘ └────┘ └────┘ └────┘
+
+When it reaches the root level, it is ready to commit the new btree!::
+
+          ┌─────────┐
+          │  root   │
+          │   PP    │
+          └─────────┘
+          ↙         ↘
+      ┌────┐       ┌────┐
+      │node│──────→│node│
+      │PP  │←──────│PP  │
+      └────┘       └────┘
+      ↙   ↘         ↙   ↘
+  ┌────┐ ┌────┐ ┌────┐ ┌────┐
+  │leaf│→│leaf│→│leaf│→│leaf│
+  │RRR │←│RRR │←│RRR │←│RRR │
+  └────┘ └────┘ └────┘ └────┘
+
+The first step to commit the new btree is to persist the btree blocks to disk
+synchronously.
+This is a little complicated because a new btree block could have been freed
+in the recent past, so the builder must use ``xfs_buf_delwri_queue_here`` to
+remove the (stale) buffer from the AIL list before it can write the new blocks
+to disk.
+Blocks are queued for IO using a delwri list and written in one large batch
+with ``xfs_buf_delwri_submit``.
+
+Once the new blocks have been persisted to disk, control returns to the
+individual repair function that called the bulk loader.
+The repair function must log the location of the new root in a transaction,
+clean up the space reservations that were made for the new btree, and reap the
+old metadata blocks:
+
+1. Commit the location of the new btree root.
+
+2. For each incore reservation:
+
+   a. Log Extent Freeing Done (EFD) items for all the space that was consumed
+      by the btree builder.  The new EFDs must point to the EFIs attached to
+      the reservation to prevent log recovery from freeing the new blocks.
+
+   b. For unclaimed portions of incore reservations, create a regular deferred
+      extent free work item to be free the unused space later in the
+      transaction chain.
+
+   c. The EFDs and EFIs logged in steps 2a and 2b must not overrun the
+      reservation of the committing transaction.
+      If the btree loading code suspects this might be about to happen, it must
+      call ``xrep_defer_finish`` to clear out the deferred work and obtain a
+      fresh transaction.
+
+3. Clear out the deferred work a second time to finish the commit and clean
+   the repair transaction.
+
+The transaction rolling in steps 2c and 3 represent a weakness in the repair
+algorithm, because a log flush and a crash before the end of the reap step can
+result in space leaking.
+Online repair functions minimize the chances of this occurring by using very
+large transactions, which each can accommodate many thousands of block freeing
+instructions.
+Repair moves on to reaping the old blocks, which will be presented in a
+subsequent :ref:`section<reaping>` after a few case studies of bulk loading.
+
+Case Study: Rebuilding the Inode Index
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild the inode index btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_inobt_rec``
+   records from the inode chunk information and a bitmap of the old inode btree
+   blocks.
+
+2. Append the records to an xfarray in inode order.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the inode btree.
+   If the free space inode btree is enabled, call it again to estimate the
+   geometry of the finobt.
+
+4. Allocate the number of blocks computed in the previous step.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+   If the free space inode btree is enabled, call it again to load the finobt.
+
+6. Commit the location of the new btree root block(s) to the AGI.
+
+7. Reap the old btree blocks using the bitmap created in step 1.
+
+Details are as follows.
+
+The inode btree maps inumbers to the ondisk location of the associated
+inode records, which means that the inode btrees can be rebuilt from the
+reverse mapping information.
+Reverse mapping records with an owner of ``XFS_RMAP_OWN_INOBT`` marks the
+location of the old inode btree blocks.
+Each reverse mapping record with an owner of ``XFS_RMAP_OWN_INODES`` marks the
+location of at least one inode cluster buffer.
+A cluster is the smallest number of ondisk inodes that can be allocated or
+freed in a single transaction; it is never smaller than 1 fs block or 4 inodes.
+
+For the space represented by each inode cluster, ensure that there are no
+records in the free space btrees nor any records in the reference count btree.
+If there are, the space metadata inconsistencies are reason enough to abort the
+operation.
+Otherwise, read each cluster buffer to check that its contents appear to be
+ondisk inodes and to decide if the file is allocated
+(``xfs_dinode.i_mode != 0``) or free (``xfs_dinode.i_mode == 0``).
+Accumulate the results of successive inode cluster buffer reads until there is
+enough information to fill a single inode chunk record, which is 64 consecutive
+numbers in the inumber keyspace.
+If the chunk is sparse, the chunk record may include holes.
+
+Once the repair function accumulates one chunk's worth of data, it calls
+``xfarray_append`` to add the inode btree record to the xfarray.
+This xfarray is walked twice during the btree creation step -- once to populate
+the inode btree with all inode chunk records, and a second time to populate the
+free inode btree with records for chunks that have free non-sparse inodes.
+The number of records for the inode btree is the number of xfarray records,
+but the record count for the free inode btree has to be computed as inode chunk
+records are stored in the xfarray.
+
+Case Study: Rebuilding the Space Reference Counts
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Reverse mapping records are used to rebuild the reference count information.
+Reference counts are required for correct operation of copy on write for shared
+file data.
+Imagine the reverse mapping entries as rectangles representing extents of
+physical blocks, and that the rectangles can be laid down to allow them to
+overlap each other.
+From the diagram below, it is apparent that a reference count record must start
+or end wherever the height of the stack changes.
+In other words, the record emission stimulus is level-triggered::
+
+                        █    ███
+              ██      █████ ████   ███        ██████
+        ██   ████     ███████████ ████     █████████
+        ████████████████████████████████ ███████████
+        ^ ^  ^^ ^^    ^ ^^ ^^^  ^^^^  ^ ^^ ^  ^     ^
+        2 1  23 21    3 43 234  2123  1 01 2  3     0
+
+The ondisk reference count btree does not store the refcount == 0 cases because
+the free space btree already records which blocks are free.
+Extents being used to stage copy-on-write operations should be the only records
+with refcount == 1.
+Single-owner file blocks aren't recorded in either the free space or the
+reference count btrees.
+
+The high level process to rebuild the reference count btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_refcount_irec``
+   records for any space having more than one reverse mapping and add them to
+   the xfarray.
+   Any records owned by ``XFS_RMAP_OWN_COW`` are also added to the xfarray
+   because these are extents allocated to stage a copy on write operation and
+   are tracked in the refcount btree.
+
+   Use any records owned by ``XFS_RMAP_OWN_REFC`` to create a bitmap of old
+   refcount btree blocks.
+
+2. Sort the records in physical extent order, putting the CoW staging extents
+   at the end of the xfarray.
+   This matches the sorting order of records in the refcount btree.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the new tree.
+
+4. Allocate the number of blocks computed in the previous step.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+
+6. Commit the location of new btree root block to the AGF.
+
+7. Reap the old btree blocks using the bitmap created in step 1.
+
+Details are as follows; the same algorithm is used by ``xfs_repair`` to
+generate refcount information from reverse mapping records.
+
+- Until the reverse mapping btree runs out of records:
+
+  - Retrieve the next record from the btree and put it in a bag.
+
+  - Collect all records with the same starting block from the btree and put
+    them in the bag.
+
+  - While the bag isn't empty:
+
+    - Among the mappings in the bag, compute the lowest block number where the
+      reference count changes.
+      This position will be either the starting block number of the next
+      unprocessed reverse mapping or the next block after the shortest mapping
+      in the bag.
+
+    - Remove all mappings from the bag that end at this position.
+
+    - Collect all reverse mappings that start at this position from the btree
+      and put them in the bag.
+
+    - If the size of the bag changed and is greater than one, create a new
+      refcount record associating the block number range that we just walked to
+      the size of the bag.
+
+The bag-like structure in this case is a type 2 xfarray as discussed in the
+:ref:`xfarray access patterns<xfarray_access_patterns>` section.
+Reverse mappings are added to the bag using ``xfarray_store_anywhere`` and
+removed via ``xfarray_unset``.
+Bag members are examined through ``xfarray_iter`` loops.
+
+Case Study: Rebuilding File Fork Mapping Indices
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+The high level process to rebuild a data/attr fork mapping btree is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_bmbt_rec``
+   records from the reverse mapping records for that inode and fork.
+   Append these records to an xfarray.
+   Compute the bitmap of the old bmap btree blocks from the ``BMBT_BLOCK``
+   records.
+
+2. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for the new tree.
+
+3. Sort the records in file offset order.
+
+4. If the extent records would fit in the inode fork immediate area, commit the
+   records to that immediate area and skip to step 8.
+
+5. Allocate the number of blocks computed in the previous step.
+
+6. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks.
+
+7. Commit the new btree root block to the inode fork immediate area.
+
+8. Reap the old btree blocks using the bitmap created in step 1.
+
+There are some complications here:
+First, it's possible to move the fork offset to adjust the sizes of the
+immediate areas if the data and attr forks are not both in BMBT format.
+Second, if there are sufficiently few fork mappings, it may be possible to use
+EXTENTS format instead of BMBT, which may require a conversion.
+Third, the incore extent map must be reloaded carefully to avoid disturbing
+any delayed allocation extents.
+
+.. _reaping:
+
+Reaping Old Metadata Blocks
+---------------------------
+
+Whenever online fsck builds a new data structure to replace one that is
+suspect, there is a question of how to find and dispose of the blocks that
+belonged to the old structure.
+The laziest method of course is not to deal with them at all, but this slowly
+leads to service degradations as space leaks out of the filesystem.
+Hopefully, someone will schedule a rebuild of the free space information to
+plug all those leaks.
+Offline repair rebuilds all space metadata after recording the usage of
+the files and directories that it decides not to clear, hence it can build new
+structures in the discovered free space and avoid the question of reaping.
+
+As part of a repair, online fsck relies heavily on the reverse mapping records
+to find space that is owned by the corresponding rmap owner yet truly free.
+Cross referencing rmap records with other rmap records is necessary because
+there may be other data structures that also think they own some of those
+blocks (e.g. crosslinked trees).
+Permitting the block allocator to hand them out again will not push the system
+towards consistency.
+
+For space metadata, the process of finding extents to dispose of generally
+follows this format:
+
+1. Create a bitmap of space used by data structures that must be preserved.
+   The space reservations used to create the new metadata can be used here if
+   the same rmap owner code is used to denote all of the objects being rebuilt.
+
+2. Survey the reverse mapping data to create a bitmap of space owned by the
+   same ``XFS_RMAP_OWN_*`` number for the metadata that is being preserved.
+
+3. Use the bitmap disunion operator to subtract (1) from (2).
+   The remaining set bits represent candidate extents that could be freed.
+   The process moves on to step 4 below.
+
+Repairs for file-based metadata such as extended attributes, directories,
+symbolic links, quota files and realtime bitmaps are performed by building a
+new structure attached to a temporary file and exchanging all mappings in the
+file forks.
+Afterward, the mappings in the old file fork are the candidate blocks for
+disposal.
+
+The process for disposing of old extents is as follows:
+
+4. For each candidate extent, count the number of reverse mapping records for
+   the first block in that extent that do not have the same rmap owner for the
+   data structure being repaired.
+
+   - If zero, the block has a single owner and can be freed.
+
+   - If not, the block is part of a crosslinked structure and must not be
+     freed.
+
+5. Starting with the next block in the extent, figure out how many more blocks
+   have the same zero/nonzero other owner status as that first block.
+
+6. If the region is crosslinked, delete the reverse mapping entry for the
+   structure being repaired and move on to the next region.
+
+7. If the region is to be freed, mark any corresponding buffers in the buffer
+   cache as stale to prevent log writeback.
+
+8. Free the region and move on.
+
+However, there is one complication to this procedure.
+Transactions are of finite size, so the reaping process must be careful to roll
+the transactions to avoid overruns.
+Overruns come from two sources:
+
+a. EFIs logged on behalf of space that is no longer occupied
+
+b. Log items for buffer invalidations
+
+This is also a window in which a crash during the reaping process can leak
+blocks.
+As stated earlier, online repair functions use very large transactions to
+minimize the chances of this occurring.
+
+Case Study: Reaping After a Regular Btree Repair
+````````````````````````````````````````````````
+
+Old reference count and inode btrees are the easiest to reap because they have
+rmap records with special owner codes: ``XFS_RMAP_OWN_REFC`` for the refcount
+btree, and ``XFS_RMAP_OWN_INOBT`` for the inode and free inode btrees.
+Creating a list of extents to reap the old btree blocks is quite simple,
+conceptually:
+
+1. Lock the relevant AGI/AGF header buffers to prevent allocation and frees.
+
+2. For each reverse mapping record with an rmap owner corresponding to the
+   metadata structure being rebuilt, set the corresponding range in a bitmap.
+
+3. Walk the current data structures that have the same rmap owner.
+   For each block visited, clear that range in the above bitmap.
+
+4. Each set bit in the bitmap represents a block that could be a block from the
+   old data structures and hence is a candidate for reaping.
+   In other words, ``(rmap_records_owned_by & ~blocks_reachable_by_walk)``
+   are the blocks that might be freeable.
+
+If it is possible to maintain the AGF lock throughout the repair (which is the
+common case), then step 2 can be performed at the same time as the reverse
+mapping record walk that creates the records for the new btree.
+
+Case Study: Rebuilding the Free Space Indices
+`````````````````````````````````````````````
+
+The high level process to rebuild the free space indices is:
+
+1. Walk the reverse mapping records to generate ``struct xfs_alloc_rec_incore``
+   records from the gaps in the reverse mapping btree.
+
+2. Append the records to an xfarray.
+
+3. Use the ``xfs_btree_bload_compute_geometry`` function to compute the number
+   of blocks needed for each new tree.
+
+4. Allocate the number of blocks computed in the previous step from the free
+   space information collected.
+
+5. Use ``xfs_btree_bload`` to write the xfarray records to btree blocks and
+   generate the internal node blocks for the free space by length index.
+   Call it again for the free space by block number index.
+
+6. Commit the locations of the new btree root blocks to the AGF.
+
+7. Reap the old btree blocks by looking for space that is not recorded by the
+   reverse mapping btree, the new free space btrees, or the AGFL.
+
+Repairing the free space btrees has three key complications over a regular
+btree repair:
+
+First, free space is not explicitly tracked in the reverse mapping records.
+Hence, the new free space records must be inferred from gaps in the physical
+space component of the keyspace of the reverse mapping btree.
+
+Second, free space repairs cannot use the common btree reservation code because
+new blocks are reserved out of the free space btrees.
+This is impossible when repairing the free space btrees themselves.
+However, repair holds the AGF buffer lock for the duration of the free space
+index reconstruction, so it can use the collected free space information to
+supply the blocks for the new free space btrees.
+It is not necessary to back each reserved extent with an EFI because the new
+free space btrees are constructed in what the ondisk filesystem thinks is
+unowned space.
+However, if reserving blocks for the new btrees from the collected free space
+information changes the number of free space records, repair must re-estimate
+the new free space btree geometry with the new record count until the
+reservation is sufficient.
+As part of committing the new btrees, repair must ensure that reverse mappings
+are created for the reserved blocks and that unused reserved blocks are
+inserted into the free space btrees.
+Deferrred rmap and freeing operations are used to ensure that this transition
+is atomic, similar to the other btree repair functions.
+
+Third, finding the blocks to reap after the repair is not overly
+straightforward.
+Blocks for the free space btrees and the reverse mapping btrees are supplied by
+the AGFL.
+Blocks put onto the AGFL have reverse mapping records with the owner
+``XFS_RMAP_OWN_AG``.
+This ownership is retained when blocks move from the AGFL into the free space
+btrees or the reverse mapping btrees.
+When repair walks reverse mapping records to synthesize free space records, it
+creates a bitmap (``ag_owner_bitmap``) of all the space claimed by
+``XFS_RMAP_OWN_AG`` records.
+The repair context maintains a second bitmap corresponding to the rmap btree
+blocks and the AGFL blocks (``rmap_agfl_bitmap``).
+When the walk is complete, the bitmap disunion operation ``(ag_owner_bitmap &
+~rmap_agfl_bitmap)`` computes the extents that are used by the old free space
+btrees.
+These blocks can then be reaped using the methods outlined above.
+
+.. _rmap_reap:
+
+Case Study: Reaping After Repairing Reverse Mapping Btrees
+``````````````````````````````````````````````````````````
+
+Old reverse mapping btrees are less difficult to reap after a repair.
+As mentioned in the previous section, blocks on the AGFL, the two free space
+btree blocks, and the reverse mapping btree blocks all have reverse mapping
+records with ``XFS_RMAP_OWN_AG`` as the owner.
+The full process of gathering reverse mapping records and building a new btree
+are described in the case study of
+:ref:`live rebuilds of rmap data <rmap_repair>`, but a crucial point from that
+discussion is that the new rmap btree will not contain any records for the old
+rmap btree, nor will the old btree blocks be tracked in the free space btrees.
+The list of candidate reaping blocks is computed by setting the bits
+corresponding to the gaps in the new rmap btree records, and then clearing the
+bits corresponding to extents in the free space btrees and the current AGFL
+blocks.
+The result ``(new_rmapbt_gaps & ~(agfl | bnobt_records))`` are reaped using the
+methods outlined above.
+
+The rest of the process of rebuildng the reverse mapping btree is discussed
+in a separate :ref:`case study<rmap_repair>`.
+
+Case Study: Rebuilding the AGFL
+```````````````````````````````
+
+The allocation group free block list (AGFL) is repaired as follows:
+
+1. Create a bitmap for all the space that the reverse mapping data claims is
+   owned by ``XFS_RMAP_OWN_AG``.
+
+2. Subtract the space used by the two free space btrees and the rmap btree.
+
+3. Subtract any space that the reverse mapping data claims is owned by any
+   other owner, to avoid re-adding crosslinked blocks to the AGFL.
+
+4. Once the AGFL is full, reap any blocks leftover.
+
+5. The next operation to fix the freelist will right-size the list.
+
+See `fs/xfs/scrub/agheader_repair.c <https://git.kernel.org/pub/scm/linux/kernel/git/torvalds/linux.git/tree/fs/xfs/scrub/agheader_repair.c>`_ for more details.
+
+Inode Record Repairs
+--------------------
+
+Inode records must be handled carefully, because they have both ondisk records
+("dinodes") and an in-memory ("cached") representation.
+There is a very high potential for cache coherency issues if online fsck is not
+careful to access the ondisk metadata *only* when the ondisk metadata is so
+badly damaged that the filesystem cannot load the in-memory representation.
+When online fsck wants to open a damaged file for scrubbing, it must use
+specialized resource acquisition functions that return either the in-memory
+representation *or* a lock on whichever object is necessary to prevent any
+update to the ondisk location.
+
+The only repairs that should be made to the ondisk inode buffers are whatever
+is necessary to get the in-core structure loaded.
+This means fixing whatever is caught by the inode cluster buffer and inode fork
+verifiers, and retrying the ``iget`` operation.
+If the second ``iget`` fails, the repair has failed.
+
+Once the in-memory representation is loaded, repair can lock the inode and can
+subject it to comprehensive checks, repairs, and optimizations.
+Most inode attributes are easy to check and constrain, or are user-controlled
+arbitrary bit patterns; these are both easy to fix.
+Dealing with the data and attr fork extent counts and the file block counts is
+more complicated, because computing the correct value requires traversing the
+forks, or if that fails, leaving the fields invalid and waiting for the fork
+fsck functions to run.
+
+Quota Record Repairs
+--------------------
+
+Similar to inodes, quota records ("dquots") also have both ondisk records and
+an in-memory representation, and hence are subject to the same cache coherency
+issues.
+Somewhat confusingly, both are known as dquots in the XFS codebase.
+
+The only repairs that should be made to the ondisk quota record buffers are
+whatever is necessary to get the in-core structure loaded.
+Once the in-memory representation is loaded, the only attributes needing
+checking are obviously bad limits and timer values.
+
+Quota usage counters are checked, repaired, and discussed separately in the
+section about :ref:`live quotacheck <quotacheck>`.
+
+.. _fscounters:
+
+Freezing to Fix Summary Counters
+--------------------------------
+
+Filesystem summary counters track availability of filesystem resources such
+as free blocks, free inodes, and allocated inodes.
+This information could be compiled by walking the free space and inode indexes,
+but this is a slow process, so XFS maintains a copy in the ondisk superblock
+that should reflect the ondisk metadata, at least when the filesystem has been
+unmounted cleanly.
+For performance reasons, XFS also maintains incore copies of those counters,
+which are key to enabling resource reservations for active transactions.
+Writer threads reserve the worst-case quantities of resources from the
+incore counter and give back whatever they don't use at commit time.
+It is therefore only necessary to serialize on the superblock when the
+superblock is being committed to disk.
+
+The lazy superblock counter feature introduced in XFS v5 took this even further
+by training log recovery to recompute the summary counters from the AG headers,
+which eliminated the need for most transactions even to touch the superblock.
+The only time XFS commits the summary counters is at filesystem unmount.
+To reduce contention even further, the incore counter is implemented as a
+percpu counter, which means that each CPU is allocated a batch of blocks from a
+global incore counter and can satisfy small allocations from the local batch.
+
+The high-performance nature of the summary counters makes it difficult for
+online fsck to check them, since there is no way to quiesce a percpu counter
+while the system is running.
+Although online fsck can read the filesystem metadata to compute the correct
+values of the summary counters, there's no way to hold the value of a percpu
+counter stable, so it's quite possible that the counter will be out of date by
+the time the walk is complete.
+Earlier versions of online scrub would return to userspace with an incomplete
+scan flag, but this is not a satisfying outcome for a system administrator.
+For repairs, the in-memory counters must be stabilized while walking the
+filesystem metadata to get an accurate reading and install it in the percpu
+counter.
+
+To satisfy this requirement, online fsck must prevent other programs in the
+system from initiating new writes to the filesystem, it must disable background
+garbage collection threads, and it must wait for existing writer programs to
+exit the kernel.
+Once that has been established, scrub can walk the AG free space indexes, the
+inode btrees, and the realtime bitmap to compute the correct value of all
+four summary counters.
+This is very similar to a filesystem freeze, though not all of the pieces are
+necessary:
+
+- The final freeze state is set one higher than ``SB_FREEZE_COMPLETE`` to
+  prevent other threads from thawing the filesystem, or other scrub threads
+  from initiating another fscounters freeze.
+
+- It does not quiesce the log.
+
+With this code in place, it is now possible to pause the filesystem for just
+long enough to check and correct the summary counters.
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| The initial implementation used the actual VFS filesystem freeze         |
+| mechanism to quiesce filesystem activity.                                |
+| With the filesystem frozen, it is possible to resolve the counter values |
+| with exact precision, but there are many problems with calling the VFS   |
+| methods directly:                                                        |
+|                                                                          |
+| - Other programs can unfreeze the filesystem without our knowledge.      |
+|   This leads to incorrect scan results and incorrect repairs.            |
+|                                                                          |
+| - Adding an extra lock to prevent others from thawing the filesystem     |
+|   required the addition of a ``->freeze_super`` function to wrap         |
+|   ``freeze_fs()``.                                                       |
+|   This in turn caused other subtle problems because it turns out that    |
+|   the VFS ``freeze_super`` and ``thaw_super`` functions can drop the     |
+|   last reference to the VFS superblock, and any subsequent access        |
+|   becomes a UAF bug!                                                     |
+|   This can happen if the filesystem is unmounted while the underlying    |
+|   block device has frozen the filesystem.                                |
+|   This problem could be solved by grabbing extra references to the       |
+|   superblock, but it felt suboptimal given the other inadequacies of     |
+|   this approach.                                                         |
+|                                                                          |
+| - The log need not be quiesced to check the summary counters, but a VFS  |
+|   freeze initiates one anyway.                                           |
+|   This adds unnecessary runtime to live fscounter fsck operations.       |
+|                                                                          |
+| - Quiescing the log means that XFS flushes the (possibly incorrect)      |
+|   counters to disk as part of cleaning the log.                          |
+|                                                                          |
+| - A bug in the VFS meant that freeze could complete even when            |
+|   sync_filesystem fails to flush the filesystem and returns an error.    |
+|   This bug was fixed in Linux 5.17.                                      |
++--------------------------------------------------------------------------+
+
+Full Filesystem Scans
+---------------------
+
+Certain types of metadata can only be checked by walking every file in the
+entire filesystem to record observations and comparing the observations against
+what's recorded on disk.
+Like every other type of online repair, repairs are made by writing those
+observations to disk in a replacement structure and committing it atomically.
+However, it is not practical to shut down the entire filesystem to examine
+hundreds of billions of files because the downtime would be excessive.
+Therefore, online fsck must build the infrastructure to manage a live scan of
+all the files in the filesystem.
+There are two questions that need to be solved to perform a live walk:
+
+- How does scrub manage the scan while it is collecting data?
+
+- How does the scan keep abreast of changes being made to the system by other
+  threads?
+
+.. _iscan:
+
+Coordinated Inode Scans
+```````````````````````
+
+In the original Unix filesystems of the 1970s, each directory entry contained
+an index number (*inumber*) which was used as an index into on ondisk array
+(*itable*) of fixed-size records (*inodes*) describing a file's attributes and
+its data block mapping.
+This system is described by J. Lions, `"inode (5659)"
+<http://www.lemis.com/grog/Documentation/Lions/>`_ in *Lions' Commentary on
+UNIX, 6th Edition*, (Dept. of Computer Science, the University of New South
+Wales, November 1977), pp. 18-2; and later by D. Ritchie and K. Thompson,
+`"Implementation of the File System"
+<https://archive.org/details/bstj57-6-1905/page/n8/mode/1up>`_, from *The UNIX
+Time-Sharing System*, (The Bell System Technical Journal, July 1978), pp.
+1913-4.
+
+XFS retains most of this design, except now inumbers are search keys over all
+the space in the data section filesystem.
+They form a continuous keyspace that can be expressed as a 64-bit integer,
+though the inodes themselves are sparsely distributed within the keyspace.
+Scans proceed in a linear fashion across the inumber keyspace, starting from
+``0x0`` and ending at ``0xFFFFFFFFFFFFFFFF``.
+Naturally, a scan through a keyspace requires a scan cursor object to track the
+scan progress.
+Because this keyspace is sparse, this cursor contains two parts.
+The first part of this scan cursor object tracks the inode that will be
+examined next; call this the examination cursor.
+Somewhat less obviously, the scan cursor object must also track which parts of
+the keyspace have already been visited, which is critical for deciding if a
+concurrent filesystem update needs to be incorporated into the scan data.
+Call this the visited inode cursor.
+
+Advancing the scan cursor is a multi-step process encapsulated in
+``xchk_iscan_iter``:
+
+1. Lock the AGI buffer of the AG containing the inode pointed to by the visited
+   inode cursor.
+   This guarantee that inodes in this AG cannot be allocated or freed while
+   advancing the cursor.
+
+2. Use the per-AG inode btree to look up the next inumber after the one that
+   was just visited, since it may not be keyspace adjacent.
+
+3. If there are no more inodes left in this AG:
+
+   a. Move the examination cursor to the point of the inumber keyspace that
+      corresponds to the start of the next AG.
+
+   b. Adjust the visited inode cursor to indicate that it has "visited" the
+      last possible inode in the current AG's inode keyspace.
+      XFS inumbers are segmented, so the cursor needs to be marked as having
+      visited the entire keyspace up to just before the start of the next AG's
+      inode keyspace.
+
+   c. Unlock the AGI and return to step 1 if there are unexamined AGs in the
+      filesystem.
+
+   d. If there are no more AGs to examine, set both cursors to the end of the
+      inumber keyspace.
+      The scan is now complete.
+
+4. Otherwise, there is at least one more inode to scan in this AG:
+
+   a. Move the examination cursor ahead to the next inode marked as allocated
+      by the inode btree.
+
+   b. Adjust the visited inode cursor to point to the inode just prior to where
+      the examination cursor is now.
+      Because the scanner holds the AGI buffer lock, no inodes could have been
+      created in the part of the inode keyspace that the visited inode cursor
+      just advanced.
+
+5. Get the incore inode for the inumber of the examination cursor.
+   By maintaining the AGI buffer lock until this point, the scanner knows that
+   it was safe to advance the examination cursor across the entire keyspace,
+   and that it has stabilized this next inode so that it cannot disappear from
+   the filesystem until the scan releases the incore inode.
+
+6. Drop the AGI lock and return the incore inode to the caller.
+
+Online fsck functions scan all files in the filesystem as follows:
+
+1. Start a scan by calling ``xchk_iscan_start``.
+
+2. Advance the scan cursor (``xchk_iscan_iter``) to get the next inode.
+   If one is provided:
+
+   a. Lock the inode to prevent updates during the scan.
+
+   b. Scan the inode.
+
+   c. While still holding the inode lock, adjust the visited inode cursor
+      (``xchk_iscan_mark_visited``) to point to this inode.
+
+   d. Unlock and release the inode.
+
+8. Call ``xchk_iscan_teardown`` to complete the scan.
+
+There are subtleties with the inode cache that complicate grabbing the incore
+inode for the caller.
+Obviously, it is an absolute requirement that the inode metadata be consistent
+enough to load it into the inode cache.
+Second, if the incore inode is stuck in some intermediate state, the scan
+coordinator must release the AGI and push the main filesystem to get the inode
+back into a loadable state.
+
+Inode Management
+````````````````
+
+In regular filesystem code, references to allocated XFS incore inodes are
+always obtained (``xfs_iget``) outside of transaction context because the
+creation of the incore context for an existing file does not require metadata
+updates.
+However, it is important to note that references to incore inodes obtained as
+part of file creation must be performed in transaction context because the
+filesystem must ensure the atomicity of the ondisk inode btree index updates
+and the initialization of the actual ondisk inode.
+
+References to incore inodes are always released (``xfs_irele``) outside of
+transaction context because there are a handful of activities that might
+require ondisk updates:
+
+- The VFS may decide to kick off writeback as part of a ``DONTCACHE`` inode
+  release.
+
+- Speculative preallocations need to be unreserved.
+
+- An unlinked file may have lost its last reference, in which case the entire
+  file must be inactivated, which involves releasing all of its resources in
+  the ondisk metadata and freeing the inode.
+
+These activities are collectively called inode inactivation.
+Inactivation has two parts -- the VFS part, which initiates writeback on all
+dirty file pages, and the XFS part, which cleans up XFS-specific information
+and frees the inode if it was unlinked.
+If the inode is unlinked (or unconnected after a file handle operation), the
+kernel drops the inode into the inactivation machinery immediately.
+
+During normal operation, resource acquisition for an update follows this order
+to avoid deadlocks:
+
+1. Inode reference (``iget``).
+
+2. Filesystem freeze protection, if repairing (``mnt_want_write_file``).
+
+3. Inode ``IOLOCK`` (VFS ``i_rwsem``) lock to control file IO.
+
+4. Inode ``MMAPLOCK`` (page cache ``invalidate_lock``) lock for operations that
+   can update page cache mappings.
+
+5. Log feature enablement.
+
+6. Transaction log space grant.
+
+7. Space on the data and realtime devices for the transaction.
+
+8. Incore dquot references, if a file is being repaired.
+   Note that they are not locked, merely acquired.
+
+9. Inode ``ILOCK`` for file metadata updates.
+
+10. AG header buffer locks / Realtime metadata inode ILOCK.
+
+11. Realtime metadata buffer locks, if applicable.
+
+12. Extent mapping btree blocks, if applicable.
+
+Resources are often released in the reverse order, though this is not required.
+However, online fsck differs from regular XFS operations because it may examine
+an object that normally is acquired in a later stage of the locking order, and
+then decide to cross-reference the object with an object that is acquired
+earlier in the order.
+The next few sections detail the specific ways in which online fsck takes care
+to avoid deadlocks.
+
+iget and irele During a Scrub
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+An inode scan performed on behalf of a scrub operation runs in transaction
+context, and possibly with resources already locked and bound to it.
+This isn't much of a problem for ``iget`` since it can operate in the context
+of an existing transaction, as long as all of the bound resources are acquired
+before the inode reference in the regular filesystem.
+
+When the VFS ``iput`` function is given a linked inode with no other
+references, it normally puts the inode on an LRU list in the hope that it can
+save time if another process re-opens the file before the system runs out
+of memory and frees it.
+Filesystem callers can short-circuit the LRU process by setting a ``DONTCACHE``
+flag on the inode to cause the kernel to try to drop the inode into the
+inactivation machinery immediately.
+
+In the past, inactivation was always done from the process that dropped the
+inode, which was a problem for scrub because scrub may already hold a
+transaction, and XFS does not support nesting transactions.
+On the other hand, if there is no scrub transaction, it is desirable to drop
+otherwise unused inodes immediately to avoid polluting caches.
+To capture these nuances, the online fsck code has a separate ``xchk_irele``
+function to set or clear the ``DONTCACHE`` flag to get the required release
+behavior.
+
+.. _ilocking:
+
+Locking Inodes
+^^^^^^^^^^^^^^
+
+In regular filesystem code, the VFS and XFS will acquire multiple IOLOCK locks
+in a well-known order: parent → child when updating the directory tree, and
+in numerical order of the addresses of their ``struct inode`` object otherwise.
+For regular files, the MMAPLOCK can be acquired after the IOLOCK to stop page
+faults.
+If two MMAPLOCKs must be acquired, they are acquired in numerical order of
+the addresses of their ``struct address_space`` objects.
+Due to the structure of existing filesystem code, IOLOCKs and MMAPLOCKs must be
+acquired before transactions are allocated.
+If two ILOCKs must be acquired, they are acquired in inumber order.
+
+Inode lock acquisition must be done carefully during a coordinated inode scan.
+Online fsck cannot abide these conventions, because for a directory tree
+scanner, the scrub process holds the IOLOCK of the file being scanned and it
+needs to take the IOLOCK of the file at the other end of the directory link.
+If the directory tree is corrupt because it contains a cycle, ``xfs_scrub``
+cannot use the regular inode locking functions and avoid becoming trapped in an
+ABBA deadlock.
+
+Solving both of these problems is straightforward -- any time online fsck
+needs to take a second lock of the same class, it uses trylock to avoid an ABBA
+deadlock.
+If the trylock fails, scrub drops all inode locks and use trylock loops to
+(re)acquire all necessary resources.
+Trylock loops enable scrub to check for pending fatal signals, which is how
+scrub avoids deadlocking the filesystem or becoming an unresponsive process.
+However, trylock loops means that online fsck must be prepared to measure the
+resource being scrubbed before and after the lock cycle to detect changes and
+react accordingly.
+
+.. _dirparent:
+
+Case Study: Finding a Directory Parent
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Consider the directory parent pointer repair code as an example.
+Online fsck must verify that the dotdot dirent of a directory points up to a
+parent directory, and that the parent directory contains exactly one dirent
+pointing down to the child directory.
+Fully validating this relationship (and repairing it if possible) requires a
+walk of every directory on the filesystem while holding the child locked, and
+while updates to the directory tree are being made.
+The coordinated inode scan provides a way to walk the filesystem without the
+possibility of missing an inode.
+The child directory is kept locked to prevent updates to the dotdot dirent, but
+if the scanner fails to lock a parent, it can drop and relock both the child
+and the prospective parent.
+If the dotdot entry changes while the directory is unlocked, then a move or
+rename operation must have changed the child's parentage, and the scan can
+exit early.
+
+.. _fshooks:
+
+Filesystem Hooks
+`````````````````
+
+The second piece of support that online fsck functions need during a full
+filesystem scan is the ability to stay informed about updates being made by
+other threads in the filesystem, since comparisons against the past are useless
+in a dynamic environment.
+Two pieces of Linux kernel infrastructure enable online fsck to monitor regular
+filesystem operations: filesystem hooks and :ref:`static keys<jump_labels>`.
+
+Filesystem hooks convey information about an ongoing filesystem operation to
+a downstream consumer.
+In this case, the downstream consumer is always an online fsck function.
+Because multiple fsck functions can run in parallel, online fsck uses the Linux
+notifier call chain facility to dispatch updates to any number of interested
+fsck processes.
+Call chains are a dynamic list, which means that they can be configured at
+run time.
+Because these hooks are private to the XFS module, the information passed along
+contains exactly what the checking function needs to update its observations.
+
+The current implementation of XFS hooks uses SRCU notifier chains to reduce the
+impact to highly threaded workloads.
+Regular blocking notifier chains use a rwsem and seem to have a much lower
+overhead for single-threaded applications.
+However, it may turn out that the combination of blocking chains and static
+keys are a more performant combination; more study is needed here.
+
+The following pieces are necessary to hook a certain point in the filesystem:
+
+- A ``struct xfs_hooks`` object must be embedded in a convenient place such as
+  a well-known incore filesystem object.
+
+- Each hook must define an action code and a structure containing more context
+  about the action.
+
+- Hook providers should provide appropriate wrapper functions and structs
+  around the ``xfs_hooks`` and ``xfs_hook`` objects to take advantage of type
+  checking to ensure correct usage.
+
+- A callsite in the regular filesystem code must be chosen to call
+  ``xfs_hooks_call`` with the action code and data structure.
+  This place should be adjacent to (and not earlier than) the place where
+  the filesystem update is committed to the transaction.
+  In general, when the filesystem calls a hook chain, it should be able to
+  handle sleeping and should not be vulnerable to memory reclaim or locking
+  recursion.
+  However, the exact requirements are very dependent on the context of the hook
+  caller and the callee.
+
+- The online fsck function should define a structure to hold scan data, a lock
+  to coordinate access to the scan data, and a ``struct xfs_hook`` object.
+  The scanner function and the regular filesystem code must acquire resources
+  in the same order; see the next section for details.
+
+- The online fsck code must contain a C function to catch the hook action code
+  and data structure.
+  If the object being updated has already been visited by the scan, then the
+  hook information must be applied to the scan data.
+
+- Prior to unlocking inodes to start the scan, online fsck must call
+  ``xfs_hooks_setup`` to initialize the ``struct xfs_hook``, and
+  ``xfs_hooks_add`` to enable the hook.
+
+- Online fsck must call ``xfs_hooks_del`` to disable the hook once the scan is
+  complete.
+
+The number of hooks should be kept to a minimum to reduce complexity.
+Static keys are used to reduce the overhead of filesystem hooks to nearly
+zero when online fsck is not running.
+
+.. _liveupdate:
+
+Live Updates During a Scan
+``````````````````````````
+
+The code paths of the online fsck scanning code and the :ref:`hooked<fshooks>`
+filesystem code look like this::
+
+            other program
+                  ↓
+            inode lock ←────────────────────┐
+                  ↓                         │
+            AG header lock                  │
+                  ↓                         │
+            filesystem function             │
+                  ↓                         │
+            notifier call chain             │    same
+                  ↓                         ├─── inode
+            scrub hook function             │    lock
+                  ↓                         │
+            scan data mutex ←──┐    same    │
+                  ↓            ├─── scan    │
+            update scan data   │    lock    │
+                  ↑            │            │
+            scan data mutex ←──┘            │
+                  ↑                         │
+            inode lock ←────────────────────┘
+                  ↑
+            scrub function
+                  ↑
+            inode scanner
+                  ↑
+            xfs_scrub
+
+These rules must be followed to ensure correct interactions between the
+checking code and the code making an update to the filesystem:
+
+- Prior to invoking the notifier call chain, the filesystem function being
+  hooked must acquire the same lock that the scrub scanning function acquires
+  to scan the inode.
+
+- The scanning function and the scrub hook function must coordinate access to
+  the scan data by acquiring a lock on the scan data.
+
+- Scrub hook function must not add the live update information to the scan
+  observations unless the inode being updated has already been scanned.
+  The scan coordinator has a helper predicate (``xchk_iscan_want_live_update``)
+  for this.
+
+- Scrub hook functions must not change the caller's state, including the
+  transaction that it is running.
+  They must not acquire any resources that might conflict with the filesystem
+  function being hooked.
+
+- The hook function can abort the inode scan to avoid breaking the other rules.
+
+The inode scan APIs are pretty simple:
+
+- ``xchk_iscan_start`` starts a scan
+
+- ``xchk_iscan_iter`` grabs a reference to the next inode in the scan or
+  returns zero if there is nothing left to scan
+
+- ``xchk_iscan_want_live_update`` to decide if an inode has already been
+  visited in the scan.
+  This is critical for hook functions to decide if they need to update the
+  in-memory scan information.
+
+- ``xchk_iscan_mark_visited`` to mark an inode as having been visited in the
+  scan
+
+- ``xchk_iscan_teardown`` to finish the scan
+
+.. _quotacheck:
+
+Case Study: Quota Counter Checking
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+It is useful to compare the mount time quotacheck code to the online repair
+quotacheck code.
+Mount time quotacheck does not have to contend with concurrent operations, so
+it does the following:
+
+1. Make sure the ondisk dquots are in good enough shape that all the incore
+   dquots will actually load, and zero the resource usage counters in the
+   ondisk buffer.
+
+2. Walk every inode in the filesystem.
+   Add each file's resource usage to the incore dquot.
+
+3. Walk each incore dquot.
+   If the incore dquot is not being flushed, add the ondisk buffer backing the
+   incore dquot to a delayed write (delwri) list.
+
+4. Write the buffer list to disk.
+
+Like most online fsck functions, online quotacheck can't write to regular
+filesystem objects until the newly collected metadata reflect all filesystem
+state.
+Therefore, online quotacheck records file resource usage to a shadow dquot
+index implemented with a sparse ``xfarray``, and only writes to the real dquots
+once the scan is complete.
+Handling transactional updates is tricky because quota resource usage updates
+are handled in phases to minimize contention on dquots:
+
+1. The inodes involved are joined and locked to a transaction.
+
+2. For each dquot attached to the file:
+
+   a. The dquot is locked.
+
+   b. A quota reservation is added to the dquot's resource usage.
+      The reservation is recorded in the transaction.
+
+   c. The dquot is unlocked.
+
+3. Changes in actual quota usage are tracked in the transaction.
+
+4. At transaction commit time, each dquot is examined again:
+
+   a. The dquot is locked again.
+
+   b. Quota usage changes are logged and unused reservation is given back to
+      the dquot.
+
+   c. The dquot is unlocked.
+
+For online quotacheck, hooks are placed in steps 2 and 4.
+The step 2 hook creates a shadow version of the transaction dquot context
+(``dqtrx``) that operates in a similar manner to the regular code.
+The step 4 hook commits the shadow ``dqtrx`` changes to the shadow dquots.
+Notice that both hooks are called with the inode locked, which is how the
+live update coordinates with the inode scanner.
+
+The quotacheck scan looks like this:
+
+1. Set up a coordinated inode scan.
+
+2. For each inode returned by the inode scan iterator:
+
+   a. Grab and lock the inode.
+
+   b. Determine that inode's resource usage (data blocks, inode counts,
+      realtime blocks) and add that to the shadow dquots for the user, group,
+      and project ids associated with the inode.
+
+   c. Unlock and release the inode.
+
+3. For each dquot in the system:
+
+   a. Grab and lock the dquot.
+
+   b. Check the dquot against the shadow dquots created by the scan and updated
+      by the live hooks.
+
+Live updates are key to being able to walk every quota record without
+needing to hold any locks for a long duration.
+If repairs are desired, the real and shadow dquots are locked and their
+resource counts are set to the values in the shadow dquot.
+
+.. _nlinks:
+
+Case Study: File Link Count Checking
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+File link count checking also uses live update hooks.
+The coordinated inode scanner is used to visit all directories on the
+filesystem, and per-file link count records are stored in a sparse ``xfarray``
+indexed by inumber.
+During the scanning phase, each entry in a directory generates observation
+data as follows:
+
+1. If the entry is a dotdot (``'..'``) entry of the root directory, the
+   directory's parent link count is bumped because the root directory's dotdot
+   entry is self referential.
+
+2. If the entry is a dotdot entry of a subdirectory, the parent's backref
+   count is bumped.
+
+3. If the entry is neither a dot nor a dotdot entry, the target file's parent
+   count is bumped.
+
+4. If the target is a subdirectory, the parent's child link count is bumped.
+
+A crucial point to understand about how the link count inode scanner interacts
+with the live update hooks is that the scan cursor tracks which *parent*
+directories have been scanned.
+In other words, the live updates ignore any update about ``A → B`` when A has
+not been scanned, even if B has been scanned.
+Furthermore, a subdirectory A with a dotdot entry pointing back to B is
+accounted as a backref counter in the shadow data for A, since child dotdot
+entries affect the parent's link count.
+Live update hooks are carefully placed in all parts of the filesystem that
+create, change, or remove directory entries, since those operations involve
+bumplink and droplink.
+
+For any file, the correct link count is the number of parents plus the number
+of child subdirectories.
+Non-directories never have children of any kind.
+The backref information is used to detect inconsistencies in the number of
+links pointing to child subdirectories and the number of dotdot entries
+pointing back.
+
+After the scan completes, the link count of each file can be checked by locking
+both the inode and the shadow data, and comparing the link counts.
+A second coordinated inode scan cursor is used for comparisons.
+Live updates are key to being able to walk every inode without needing to hold
+any locks between inodes.
+If repairs are desired, the inode's link count is set to the value in the
+shadow information.
+If no parents are found, the file must be :ref:`reparented <orphanage>` to the
+orphanage to prevent the file from being lost forever.
+
+.. _rmap_repair:
+
+Case Study: Rebuilding Reverse Mapping Records
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Most repair functions follow the same pattern: lock filesystem resources,
+walk the surviving ondisk metadata looking for replacement metadata records,
+and use an :ref:`in-memory array <xfarray>` to store the gathered observations.
+The primary advantage of this approach is the simplicity and modularity of the
+repair code -- code and data are entirely contained within the scrub module,
+do not require hooks in the main filesystem, and are usually the most efficient
+in memory use.
+A secondary advantage of this repair approach is atomicity -- once the kernel
+decides a structure is corrupt, no other threads can access the metadata until
+the kernel finishes repairing and revalidating the metadata.
+
+For repairs going on within a shard of the filesystem, these advantages
+outweigh the delays inherent in locking the shard while repairing parts of the
+shard.
+Unfortunately, repairs to the reverse mapping btree cannot use the "standard"
+btree repair strategy because it must scan every space mapping of every fork of
+every file in the filesystem, and the filesystem cannot stop.
+Therefore, rmap repair foregoes atomicity between scrub and repair.
+It combines a :ref:`coordinated inode scanner <iscan>`, :ref:`live update hooks
+<liveupdate>`, and an :ref:`in-memory rmap btree <xfbtree>` to complete the
+scan for reverse mapping records.
+
+1. Set up an xfbtree to stage rmap records.
+
+2. While holding the locks on the AGI and AGF buffers acquired during the
+   scrub, generate reverse mappings for all AG metadata: inodes, btrees, CoW
+   staging extents, and the internal log.
+
+3. Set up an inode scanner.
+
+4. Hook into rmap updates for the AG being repaired so that the live scan data
+   can receive updates to the rmap btree from the rest of the filesystem during
+   the file scan.
+
+5. For each space mapping found in either fork of each file scanned,
+   decide if the mapping matches the AG of interest.
+   If so:
+
+   a. Create a btree cursor for the in-memory btree.
+
+   b. Use the rmap code to add the record to the in-memory btree.
+
+   c. Use the :ref:`special commit function <xfbtree_commit>` to write the
+      xfbtree changes to the xfile.
+
+6. For each live update received via the hook, decide if the owner has already
+   been scanned.
+   If so, apply the live update into the scan data:
+
+   a. Create a btree cursor for the in-memory btree.
+
+   b. Replay the operation into the in-memory btree.
+
+   c. Use the :ref:`special commit function <xfbtree_commit>` to write the
+      xfbtree changes to the xfile.
+      This is performed with an empty transaction to avoid changing the
+      caller's state.
+
+7. When the inode scan finishes, create a new scrub transaction and relock the
+   two AG headers.
+
+8. Compute the new btree geometry using the number of rmap records in the
+   shadow btree, like all other btree rebuilding functions.
+
+9. Allocate the number of blocks computed in the previous step.
+
+10. Perform the usual btree bulk loading and commit to install the new rmap
+    btree.
+
+11. Reap the old rmap btree blocks as discussed in the case study about how
+    to :ref:`reap after rmap btree repair <rmap_reap>`.
+
+12. Free the xfbtree now that it not needed.
+
+Staging Repairs with Temporary Files on Disk
+--------------------------------------------
+
+XFS stores a substantial amount of metadata in file forks: directories,
+extended attributes, symbolic link targets, free space bitmaps and summary
+information for the realtime volume, and quota records.
+File forks map 64-bit logical file fork space extents to physical storage space
+extents, similar to how a memory management unit maps 64-bit virtual addresses
+to physical memory addresses.
+Therefore, file-based tree structures (such as directories and extended
+attributes) use blocks mapped in the file fork offset address space that point
+to other blocks mapped within that same address space, and file-based linear
+structures (such as bitmaps and quota records) compute array element offsets in
+the file fork offset address space.
+
+Because file forks can consume as much space as the entire filesystem, repairs
+cannot be staged in memory, even when a paging scheme is available.
+Therefore, online repair of file-based metadata createas a temporary file in
+the XFS filesystem, writes a new structure at the correct offsets into the
+temporary file, and atomically exchanges all file fork mappings (and hence the
+fork contents) to commit the repair.
+Once the repair is complete, the old fork can be reaped as necessary; if the
+system goes down during the reap, the iunlink code will delete the blocks
+during log recovery.
+
+**Note**: All space usage and inode indices in the filesystem *must* be
+consistent to use a temporary file safely!
+This dependency is the reason why online repair can only use pageable kernel
+memory to stage ondisk space usage information.
+
+Exchanging metadata file mappings with a temporary file requires the owner
+field of the block headers to match the file being repaired and not the
+temporary file.
+The directory, extended attribute, and symbolic link functions were all
+modified to allow callers to specify owner numbers explicitly.
+
+There is a downside to the reaping process -- if the system crashes during the
+reap phase and the fork extents are crosslinked, the iunlink processing will
+fail because freeing space will find the extra reverse mappings and abort.
+
+Temporary files created for repair are similar to ``O_TMPFILE`` files created
+by userspace.
+They are not linked into a directory and the entire file will be reaped when
+the last reference to the file is lost.
+The key differences are that these files must have no access permission outside
+the kernel at all, they must be specially marked to prevent them from being
+opened by handle, and they must never be linked into the directory tree.
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| In the initial iteration of file metadata repair, the damaged metadata   |
+| blocks would be scanned for salvageable data; the extents in the file    |
+| fork would be reaped; and then a new structure would be built in its     |
+| place.                                                                   |
+| This strategy did not survive the introduction of the atomic repair      |
+| requirement expressed earlier in this document.                          |
+|                                                                          |
+| The second iteration explored building a second structure at a high      |
+| offset in the fork from the salvage data, reaping the old extents, and   |
+| using a ``COLLAPSE_RANGE`` operation to slide the new extents into       |
+| place.                                                                   |
+|                                                                          |
+| This had many drawbacks:                                                 |
+|                                                                          |
+| - Array structures are linearly addressed, and the regular filesystem    |
+|   codebase does not have the concept of a linear offset that could be    |
+|   applied to the record offset computation to build an alternate copy.   |
+|                                                                          |
+| - Extended attributes are allowed to use the entire attr fork offset     |
+|   address space.                                                         |
+|                                                                          |
+| - Even if repair could build an alternate copy of a data structure in a  |
+|   different part of the fork address space, the atomic repair commit     |
+|   requirement means that online repair would have to be able to perform  |
+|   a log assisted ``COLLAPSE_RANGE`` operation to ensure that the old     |
+|   structure was completely replaced.                                     |
+|                                                                          |
+| - A crash after construction of the secondary tree but before the range  |
+|   collapse would leave unreachable blocks in the file fork.              |
+|   This would likely confuse things further.                              |
+|                                                                          |
+| - Reaping blocks after a repair is not a simple operation, and           |
+|   initiating a reap operation from a restarted range collapse operation  |
+|   during log recovery is daunting.                                       |
+|                                                                          |
+| - Directory entry blocks and quota records record the file fork offset   |
+|   in the header area of each block.                                      |
+|   An atomic range collapse operation would have to rewrite this part of  |
+|   each block header.                                                     |
+|   Rewriting a single field in block headers is not a huge problem, but   |
+|   it's something to be aware of.                                         |
+|                                                                          |
+| - Each block in a directory or extended attributes btree index contains  |
+|   sibling and child block pointers.                                      |
+|   Were the atomic commit to use a range collapse operation, each block   |
+|   would have to be rewritten very carefully to preserve the graph        |
+|   structure.                                                             |
+|   Doing this as part of a range collapse means rewriting a large number  |
+|   of blocks repeatedly, which is not conducive to quick repairs.         |
+|                                                                          |
+| This lead to the introduction of temporary file staging.                 |
++--------------------------------------------------------------------------+
+
+Using a Temporary File
+``````````````````````
+
+Online repair code should use the ``xrep_tempfile_create`` function to create a
+temporary file inside the filesystem.
+This allocates an inode, marks the in-core inode private, and attaches it to
+the scrub context.
+These files are hidden from userspace, may not be added to the directory tree,
+and must be kept private.
+
+Temporary files only use two inode locks: the IOLOCK and the ILOCK.
+The MMAPLOCK is not needed here, because there must not be page faults from
+userspace for data fork blocks.
+The usage patterns of these two locks are the same as for any other XFS file --
+access to file data are controlled via the IOLOCK, and access to file metadata
+are controlled via the ILOCK.
+Locking helpers are provided so that the temporary file and its lock state can
+be cleaned up by the scrub context.
+To comply with the nested locking strategy laid out in the :ref:`inode
+locking<ilocking>` section, it is recommended that scrub functions use the
+xrep_tempfile_ilock*_nowait lock helpers.
+
+Data can be written to a temporary file by two means:
+
+1. ``xrep_tempfile_copyin`` can be used to set the contents of a regular
+   temporary file from an xfile.
+
+2. The regular directory, symbolic link, and extended attribute functions can
+   be used to write to the temporary file.
+
+Once a good copy of a data file has been constructed in a temporary file, it
+must be conveyed to the file being repaired, which is the topic of the next
+section.
+
+Logged File Content Exchanges
+-----------------------------
+
+Once repair builds a temporary file with a new data structure written into
+it, it must commit the new changes into the existing file.
+It is not possible to swap the inumbers of two files, so instead the new
+metadata must replace the old.
+This suggests the need for the ability to swap extents, but the existing extent
+swapping code used by the file defragmenting tool ``xfs_fsr`` is not sufficient
+for online repair because:
+
+a. When the reverse-mapping btree is enabled, the swap code must keep the
+   reverse mapping information up to date with every exchange of mappings.
+   Therefore, it can only exchange one mapping per transaction, and each
+   transaction is independent.
+
+b. Reverse-mapping is critical for the operation of online fsck, so the old
+   defragmentation code (which swapped entire extent forks in a single
+   operation) is not useful here.
+
+c. Defragmentation is assumed to occur between two files with identical
+   contents.
+   For this use case, an incomplete exchange will not result in a user-visible
+   change in file contents, even if the operation is interrupted.
+
+d. Online repair needs to swap the contents of two files that are by definition
+   *not* identical.
+   For directory and xattr repairs, the user-visible contents might be the
+   same, but the contents of individual blocks may be very different.
+
+e. Old blocks in the file may be cross-linked with another structure and must
+   not reappear if the system goes down mid-repair.
+
+These problems are overcome by creating a new deferred operation and a new type
+of log intent item to track the progress of an operation to exchange two file
+ranges.
+The new exchange operation type chains together the same transactions used by
+the reverse-mapping extent swap code, but records intermedia progress in the
+log so that operations can be restarted after a crash.
+This new functionality is called the file contents exchange (xfs_exchrange)
+code.
+The underlying implementation exchanges file fork mappings (xfs_exchmaps).
+The new log item records the progress of the exchange to ensure that once an
+exchange begins, it will always run to completion, even there are
+interruptions.
+The new ``XFS_SB_FEAT_INCOMPAT_EXCHRANGE`` incompatible feature flag
+in the superblock protects these new log item records from being replayed on
+old kernels.
+
++--------------------------------------------------------------------------+
+| **Sidebar: Using Log-Incompatible Feature Flags**                        |
++--------------------------------------------------------------------------+
+| Starting with XFS v5, the superblock contains a                          |
+| ``sb_features_log_incompat`` field to indicate that the log contains     |
+| records that might not readable by all kernels that could mount this     |
+| filesystem.                                                              |
+| In short, log incompat features protect the log contents against kernels |
+| that will not understand the contents.                                   |
+| Unlike the other superblock feature bits, log incompat bits are          |
+| ephemeral because an empty (clean) log does not need protection.         |
+| The log cleans itself after its contents have been committed into the    |
+| filesystem, either as part of an unmount or because the system is        |
+| otherwise idle.                                                          |
+| Because upper level code can be working on a transaction at the same     |
+| time that the log cleans itself, it is necessary for upper level code to |
+| communicate to the log when it is going to use a log incompatible        |
+| feature.                                                                 |
+|                                                                          |
+| The log coordinates access to incompatible features through the use of   |
+| one ``struct rw_semaphore`` for each feature.                            |
+| The log cleaning code tries to take this rwsem in exclusive mode to      |
+| clear the bit; if the lock attempt fails, the feature bit remains set.   |
+| The code supporting a log incompat feature should create wrapper         |
+| functions to obtain the log feature and call                             |
+| ``xfs_add_incompat_log_feature`` to set the feature bits in the primary  |
+| superblock.                                                              |
+| The superblock update is performed transactionally, so the wrapper to    |
+| obtain log assistance must be called just prior to the creation of the   |
+| transaction that uses the functionality.                                 |
+| For a file operation, this step must happen after taking the IOLOCK      |
+| and the MMAPLOCK, but before allocating the transaction.                 |
+| When the transaction is complete, the ``xlog_drop_incompat_feat``        |
+| function is called to release the feature.                               |
+| The feature bit will not be cleared from the superblock until the log    |
+| becomes clean.                                                           |
+|                                                                          |
+| Log-assisted extended attribute updates and file content exchanges bothe |
+| use log incompat features and provide convenience wrappers around the    |
+| functionality.                                                           |
++--------------------------------------------------------------------------+
+
+Mechanics of a Logged File Content Exchange
+```````````````````````````````````````````
+
+Exchanging contents between file forks is a complex task.
+The goal is to exchange all file fork mappings between two file fork offset
+ranges.
+There are likely to be many extent mappings in each fork, and the edges of
+the mappings aren't necessarily aligned.
+Furthermore, there may be other updates that need to happen after the exchange,
+such as exchanging file sizes, inode flags, or conversion of fork data to local
+format.
+This is roughly the format of the new deferred exchange-mapping work item:
+
+.. code-block:: c
+
+	struct xfs_exchmaps_intent {
+	    /* Inodes participating in the operation. */
+	    struct xfs_inode    *xmi_ip1;
+	    struct xfs_inode    *xmi_ip2;
+
+	    /* File offset range information. */
+	    xfs_fileoff_t       xmi_startoff1;
+	    xfs_fileoff_t       xmi_startoff2;
+	    xfs_filblks_t       xmi_blockcount;
+
+	    /* Set these file sizes after the operation, unless negative. */
+	    xfs_fsize_t         xmi_isize1;
+	    xfs_fsize_t         xmi_isize2;
+
+	    /* XFS_EXCHMAPS_* log operation flags */
+	    uint64_t            xmi_flags;
+	};
+
+The new log intent item contains enough information to track two logical fork
+offset ranges: ``(inode1, startoff1, blockcount)`` and ``(inode2, startoff2,
+blockcount)``.
+Each step of an exchange operation exchanges the largest file range mapping
+possible from one file to the other.
+After each step in the exchange operation, the two startoff fields are
+incremented and the blockcount field is decremented to reflect the progress
+made.
+The flags field captures behavioral parameters such as exchanging attr fork
+mappings instead of the data fork and other work to be done after the exchange.
+The two isize fields are used to exchange the file sizes at the end of the
+operation if the file data fork is the target of the operation.
+
+When the exchange is initiated, the sequence of operations is as follows:
+
+1. Create a deferred work item for the file mapping exchange.
+   At the start, it should contain the entirety of the file block ranges to be
+   exchanged.
+
+2. Call ``xfs_defer_finish`` to process the exchange.
+   This is encapsulated in ``xrep_tempexch_contents`` for scrub operations.
+   This will log an extent swap intent item to the transaction for the deferred
+   mapping exchange work item.
+
+3. Until ``xmi_blockcount`` of the deferred mapping exchange work item is zero,
+
+   a. Read the block maps of both file ranges starting at ``xmi_startoff1`` and
+      ``xmi_startoff2``, respectively, and compute the longest extent that can
+      be exchanged in a single step.
+      This is the minimum of the two ``br_blockcount`` s in the mappings.
+      Keep advancing through the file forks until at least one of the mappings
+      contains written blocks.
+      Mutual holes, unwritten extents, and extent mappings to the same physical
+      space are not exchanged.
+
+      For the next few steps, this document will refer to the mapping that came
+      from file 1 as "map1", and the mapping that came from file 2 as "map2".
+
+   b. Create a deferred block mapping update to unmap map1 from file 1.
+
+   c. Create a deferred block mapping update to unmap map2 from file 2.
+
+   d. Create a deferred block mapping update to map map1 into file 2.
+
+   e. Create a deferred block mapping update to map map2 into file 1.
+
+   f. Log the block, quota, and extent count updates for both files.
+
+   g. Extend the ondisk size of either file if necessary.
+
+   h. Log a mapping exchange done log item for th mapping exchange intent log
+      item that was read at the start of step 3.
+
+   i. Compute the amount of file range that has just been covered.
+      This quantity is ``(map1.br_startoff + map1.br_blockcount -
+      xmi_startoff1)``, because step 3a could have skipped holes.
+
+   j. Increase the starting offsets of ``xmi_startoff1`` and ``xmi_startoff2``
+      by the number of blocks computed in the previous step, and decrease
+      ``xmi_blockcount`` by the same quantity.
+      This advances the cursor.
+
+   k. Log a new mapping exchange intent log item reflecting the advanced state
+      of the work item.
+
+   l. Return the proper error code (EAGAIN) to the deferred operation manager
+      to inform it that there is more work to be done.
+      The operation manager completes the deferred work in steps 3b-3e before
+      moving back to the start of step 3.
+
+4. Perform any post-processing.
+   This will be discussed in more detail in subsequent sections.
+
+If the filesystem goes down in the middle of an operation, log recovery will
+find the most recent unfinished mapping exchange log intent item and restart
+from there.
+This is how atomic file mapping exchanges guarantees that an outside observer
+will either see the old broken structure or the new one, and never a mismash of
+both.
+
+Preparation for File Content Exchanges
+``````````````````````````````````````
+
+There are a few things that need to be taken care of before initiating an
+atomic file mapping exchange operation.
+First, regular files require the page cache to be flushed to disk before the
+operation begins, and directio writes to be quiesced.
+Like any filesystem operation, file mapping exchanges must determine the
+maximum amount of disk space and quota that can be consumed on behalf of both
+files in the operation, and reserve that quantity of resources to avoid an
+unrecoverable out of space failure once it starts dirtying metadata.
+The preparation step scans the ranges of both files to estimate:
+
+- Data device blocks needed to handle the repeated updates to the fork
+  mappings.
+- Change in data and realtime block counts for both files.
+- Increase in quota usage for both files, if the two files do not share the
+  same set of quota ids.
+- The number of extent mappings that will be added to each file.
+- Whether or not there are partially written realtime extents.
+  User programs must never be able to access a realtime file extent that maps
+  to different extents on the realtime volume, which could happen if the
+  operation fails to run to completion.
+
+The need for precise estimation increases the run time of the exchange
+operation, but it is very important to maintain correct accounting.
+The filesystem must not run completely out of free space, nor can the mapping
+exchange ever add more extent mappings to a fork than it can support.
+Regular users are required to abide the quota limits, though metadata repairs
+may exceed quota to resolve inconsistent metadata elsewhere.
+
+Special Features for Exchanging Metadata File Contents
+``````````````````````````````````````````````````````
+
+Extended attributes, symbolic links, and directories can set the fork format to
+"local" and treat the fork as a literal area for data storage.
+Metadata repairs must take extra steps to support these cases:
+
+- If both forks are in local format and the fork areas are large enough, the
+  exchange is performed by copying the incore fork contents, logging both
+  forks, and committing.
+  The atomic file mapping exchange mechanism is not necessary, since this can
+  be done with a single transaction.
+
+- If both forks map blocks, then the regular atomic file mapping exchange is
+  used.
+
+- Otherwise, only one fork is in local format.
+  The contents of the local format fork are converted to a block to perform the
+  exchange.
+  The conversion to block format must be done in the same transaction that
+  logs the initial mapping exchange intent log item.
+  The regular atomic mapping exchange is used to exchange the metadata file
+  mappings.
+  Special flags are set on the exchange operation so that the transaction can
+  be rolled one more time to convert the second file's fork back to local
+  format so that the second file will be ready to go as soon as the ILOCK is
+  dropped.
+
+Extended attributes and directories stamp the owning inode into every block,
+but the buffer verifiers do not actually check the inode number!
+Although there is no verification, it is still important to maintain
+referential integrity, so prior to performing the mapping exchange, online
+repair builds every block in the new data structure with the owner field of the
+file being repaired.
+
+After a successful exchange operation, the repair operation must reap the old
+fork blocks by processing each fork mapping through the standard :ref:`file
+extent reaping <reaping>` mechanism that is done post-repair.
+If the filesystem should go down during the reap part of the repair, the
+iunlink processing at the end of recovery will free both the temporary file and
+whatever blocks were not reaped.
+However, this iunlink processing omits the cross-link detection of online
+repair, and is not completely foolproof.
+
+Exchanging Temporary File Contents
+``````````````````````````````````
+
+To repair a metadata file, online repair proceeds as follows:
+
+1. Create a temporary repair file.
+
+2. Use the staging data to write out new contents into the temporary repair
+   file.
+   The same fork must be written to as is being repaired.
+
+3. Commit the scrub transaction, since the exchange resource estimation step
+   must be completed before transaction reservations are made.
+
+4. Call ``xrep_tempexch_trans_alloc`` to allocate a new scrub transaction with
+   the appropriate resource reservations, locks, and fill out a ``struct
+   xfs_exchmaps_req`` with the details of the exchange operation.
+
+5. Call ``xrep_tempexch_contents`` to exchange the contents.
+
+6. Commit the transaction to complete the repair.
+
+.. _rtsummary:
+
+Case Study: Repairing the Realtime Summary File
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+In the "realtime" section of an XFS filesystem, free space is tracked via a
+bitmap, similar to Unix FFS.
+Each bit in the bitmap represents one realtime extent, which is a multiple of
+the filesystem block size between 4KiB and 1GiB in size.
+The realtime summary file indexes the number of free extents of a given size to
+the offset of the block within the realtime free space bitmap where those free
+extents begin.
+In other words, the summary file helps the allocator find free extents by
+length, similar to what the free space by count (cntbt) btree does for the data
+section.
+
+The summary file itself is a flat file (with no block headers or checksums!)
+partitioned into ``log2(total rt extents)`` sections containing enough 32-bit
+counters to match the number of blocks in the rt bitmap.
+Each counter records the number of free extents that start in that bitmap block
+and can satisfy a power-of-two allocation request.
+
+To check the summary file against the bitmap:
+
+1. Take the ILOCK of both the realtime bitmap and summary files.
+
+2. For each free space extent recorded in the bitmap:
+
+   a. Compute the position in the summary file that contains a counter that
+      represents this free extent.
+
+   b. Read the counter from the xfile.
+
+   c. Increment it, and write it back to the xfile.
+
+3. Compare the contents of the xfile against the ondisk file.
+
+To repair the summary file, write the xfile contents into the temporary file
+and use atomic mapping exchange to commit the new contents.
+The temporary file is then reaped.
+
+Case Study: Salvaging Extended Attributes
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+In XFS, extended attributes are implemented as a namespaced name-value store.
+Values are limited in size to 64KiB, but there is no limit in the number of
+names.
+The attribute fork is unpartitioned, which means that the root of the attribute
+structure is always in logical block zero, but attribute leaf blocks, dabtree
+index blocks, and remote value blocks are intermixed.
+Attribute leaf blocks contain variable-sized records that associate
+user-provided names with the user-provided values.
+Values larger than a block are allocated separate extents and written there.
+If the leaf information expands beyond a single block, a directory/attribute
+btree (``dabtree``) is created to map hashes of attribute names to entries
+for fast lookup.
+
+Salvaging extended attributes is done as follows:
+
+1. Walk the attr fork mappings of the file being repaired to find the attribute
+   leaf blocks.
+   When one is found,
+
+   a. Walk the attr leaf block to find candidate keys.
+      When one is found,
+
+      1. Check the name for problems, and ignore the name if there are.
+
+      2. Retrieve the value.
+         If that succeeds, add the name and value to the staging xfarray and
+         xfblob.
+
+2. If the memory usage of the xfarray and xfblob exceed a certain amount of
+   memory or there are no more attr fork blocks to examine, unlock the file and
+   add the staged extended attributes to the temporary file.
+
+3. Use atomic file mapping exchange to exchange the new and old extended
+   attribute structures.
+   The old attribute blocks are now attached to the temporary file.
+
+4. Reap the temporary file.
+
+Fixing Directories
+------------------
+
+Fixing directories is difficult with currently available filesystem features,
+since directory entries are not redundant.
+The offline repair tool scans all inodes to find files with nonzero link count,
+and then it scans all directories to establish parentage of those linked files.
+Damaged files and directories are zapped, and files with no parent are
+moved to the ``/lost+found`` directory.
+It does not try to salvage anything.
+
+The best that online repair can do at this time is to read directory data
+blocks and salvage any dirents that look plausible, correct link counts, and
+move orphans back into the directory tree.
+The salvage process is discussed in the case study at the end of this section.
+The :ref:`file link count fsck <nlinks>` code takes care of fixing link counts
+and moving orphans to the ``/lost+found`` directory.
+
+Case Study: Salvaging Directories
+`````````````````````````````````
+
+Unlike extended attributes, directory blocks are all the same size, so
+salvaging directories is straightforward:
+
+1. Find the parent of the directory.
+   If the dotdot entry is not unreadable, try to confirm that the alleged
+   parent has a child entry pointing back to the directory being repaired.
+   Otherwise, walk the filesystem to find it.
+
+2. Walk the first partition of data fork of the directory to find the directory
+   entry data blocks.
+   When one is found,
+
+   a. Walk the directory data block to find candidate entries.
+      When an entry is found:
+
+      i. Check the name for problems, and ignore the name if there are.
+
+      ii. Retrieve the inumber and grab the inode.
+          If that succeeds, add the name, inode number, and file type to the
+          staging xfarray and xblob.
+
+3. If the memory usage of the xfarray and xfblob exceed a certain amount of
+   memory or there are no more directory data blocks to examine, unlock the
+   directory and add the staged dirents into the temporary directory.
+   Truncate the staging files.
+
+4. Use atomic file mapping exchange to exchange the new and old directory
+   structures.
+   The old directory blocks are now attached to the temporary file.
+
+5. Reap the temporary file.
+
+**Future Work Question**: Should repair revalidate the dentry cache when
+rebuilding a directory?
+
+*Answer*: Yes, it should.
+
+In theory it is necessary to scan all dentry cache entries for a directory to
+ensure that one of the following apply:
+
+1. The cached dentry reflects an ondisk dirent in the new directory.
+
+2. The cached dentry no longer has a corresponding ondisk dirent in the new
+   directory and the dentry can be purged from the cache.
+
+3. The cached dentry no longer has an ondisk dirent but the dentry cannot be
+   purged.
+   This is the problem case.
+
+Unfortunately, the current dentry cache design doesn't provide a means to walk
+every child dentry of a specific directory, which makes this a hard problem.
+There is no known solution.
+
+Parent Pointers
+```````````````
+
+A parent pointer is a piece of file metadata that enables a user to locate the
+file's parent directory without having to traverse the directory tree from the
+root.
+Without them, reconstruction of directory trees is hindered in much the same
+way that the historic lack of reverse space mapping information once hindered
+reconstruction of filesystem space metadata.
+The parent pointer feature, however, makes total directory reconstruction
+possible.
+
+XFS parent pointers contain the information needed to identify the
+corresponding directory entry in the parent directory.
+In other words, child files use extended attributes to store pointers to
+parents in the form ``(dirent_name) → (parent_inum, parent_gen)``.
+The directory checking process can be strengthened to ensure that the target of
+each dirent also contains a parent pointer pointing back to the dirent.
+Likewise, each parent pointer can be checked by ensuring that the target of
+each parent pointer is a directory and that it contains a dirent matching
+the parent pointer.
+Both online and offline repair can use this strategy.
+
++--------------------------------------------------------------------------+
+| **Historical Sidebar**:                                                  |
++--------------------------------------------------------------------------+
+| Directory parent pointers were first proposed as an XFS feature more     |
+| than a decade ago by SGI.                                                |
+| Each link from a parent directory to a child file is mirrored with an    |
+| extended attribute in the child that could be used to identify the       |
+| parent directory.                                                        |
+| Unfortunately, this early implementation had major shortcomings and was  |
+| never merged into Linux XFS:                                             |
+|                                                                          |
+| 1. The XFS codebase of the late 2000s did not have the infrastructure to |
+|    enforce strong referential integrity in the directory tree.           |
+|    It did not guarantee that a change in a forward link would always be  |
+|    followed up with the corresponding change to the reverse links.       |
+|                                                                          |
+| 2. Referential integrity was not integrated into offline repair.         |
+|    Checking and repairs were performed on mounted filesystems without    |
+|    taking any kernel or inode locks to coordinate access.                |
+|    It is not clear how this actually worked properly.                    |
+|                                                                          |
+| 3. The extended attribute did not record the name of the directory entry |
+|    in the parent, so the SGI parent pointer implementation cannot be     |
+|    used to reconnect the directory tree.                                 |
+|                                                                          |
+| 4. Extended attribute forks only support 65,536 extents, which means     |
+|    that parent pointer attribute creation is likely to fail at some      |
+|    point before the maximum file link count is achieved.                 |
+|                                                                          |
+| The original parent pointer design was too unstable for something like   |
+| a file system repair to depend on.                                       |
+| Allison Henderson, Chandan Babu, and Catherine Hoang are working on a    |
+| second implementation that solves all shortcomings of the first.         |
+| During 2022, Allison introduced log intent items to track physical       |
+| manipulations of the extended attribute structures.                      |
+| This solves the referential integrity problem by making it possible to   |
+| commit a dirent update and a parent pointer update in the same           |
+| transaction.                                                             |
+| Chandan increased the maximum extent counts of both data and attribute   |
+| forks, thereby ensuring that the extended attribute structure can grow   |
+| to handle the maximum hardlink count of any file.                        |
+|                                                                          |
+| For this second effort, the ondisk parent pointer format as originally   |
+| proposed was ``(parent_inum, parent_gen, dirent_pos) → (dirent_name)``.  |
+| The format was changed during development to eliminate the requirement   |
+| of repair tools needing to ensure that the ``dirent_pos`` field always   |
+| matched when reconstructing a directory.                                 |
+|                                                                          |
+| There were a few other ways to have solved that problem:                 |
+|                                                                          |
+| 1. The field could be designated advisory, since the other three values  |
+|    are sufficient to find the entry in the parent.                       |
+|    However, this makes indexed key lookup impossible while repairs are   |
+|    ongoing.                                                              |
+|                                                                          |
+| 2. We could allow creating directory entries at specified offsets, which |
+|    solves the referential integrity problem but runs the risk that       |
+|    dirent creation will fail due to conflicts with the free space in the |
+|    directory.                                                            |
+|                                                                          |
+|    These conflicts could be resolved by appending the directory entry    |
+|    and amending the xattr code to support updating an xattr key and      |
+|    reindexing the dabtree, though this would have to be performed with   |
+|    the parent directory still locked.                                    |
+|                                                                          |
+| 3. Same as above, but remove the old parent pointer entry and add a new  |
+|    one atomically.                                                       |
+|                                                                          |
+| 4. Change the ondisk xattr format to                                     |
+|    ``(parent_inum, name) → (parent_gen)``, which would provide the attr  |
+|    name uniqueness that we require, without forcing repair code to       |
+|    update the dirent position.                                           |
+|    Unfortunately, this requires changes to the xattr code to support     |
+|    attr names as long as 263 bytes.                                      |
+|                                                                          |
+| 5. Change the ondisk xattr format to ``(parent_inum, hash(name)) →       |
+|    (name, parent_gen)``.                                                 |
+|    If the hash is sufficiently resistant to collisions (e.g. sha256)     |
+|    then this should provide the attr name uniqueness that we require.    |
+|    Names shorter than 247 bytes could be stored directly.                |
+|                                                                          |
+| 6. Change the ondisk xattr format to ``(dirent_name) → (parent_ino,      |
+|    parent_gen)``.  This format doesn't require any of the complicated    |
+|    nested name hashing of the previous suggestions.  However, it was     |
+|    discovered that multiple hardlinks to the same inode with the same    |
+|    filename caused performance problems with hashed xattr lookups, so    |
+|    the parent inumber is now xor'd into the hash index.                  |
+|                                                                          |
+| In the end, it was decided that solution #6 was the most compact and the |
+| most performant.  A new hash function was designed for parent pointers.  |
++--------------------------------------------------------------------------+
+
+
+Case Study: Repairing Directories with Parent Pointers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Directory rebuilding uses a :ref:`coordinated inode scan <iscan>` and
+a :ref:`directory entry live update hook <liveupdate>` as follows:
+
+1. Set up a temporary directory for generating the new directory structure,
+   an xfblob for storing entry names, and an xfarray for stashing the fixed
+   size fields involved in a directory update: ``(child inumber, add vs.
+   remove, name cookie, ftype)``.
+
+2. Set up an inode scanner and hook into the directory entry code to receive
+   updates on directory operations.
+
+3. For each parent pointer found in each file scanned, decide if the parent
+   pointer references the directory of interest.
+   If so:
+
+   a. Stash the parent pointer name and an addname entry for this dirent in the
+      xfblob and xfarray, respectively.
+
+   b. When finished scanning that file or the kernel memory consumption exceeds
+      a threshold, flush the stashed updates to the temporary directory.
+
+4. For each live directory update received via the hook, decide if the child
+   has already been scanned.
+   If so:
+
+   a. Stash the parent pointer name an addname or removename entry for this
+      dirent update in the xfblob and xfarray for later.
+      We cannot write directly to the temporary directory because hook
+      functions are not allowed to modify filesystem metadata.
+      Instead, we stash updates in the xfarray and rely on the scanner thread
+      to apply the stashed updates to the temporary directory.
+
+5. When the scan is complete, replay any stashed entries in the xfarray.
+
+6. When the scan is complete, atomically exchange the contents of the temporary
+   directory and the directory being repaired.
+   The temporary directory now contains the damaged directory structure.
+
+7. Reap the temporary directory.
+
+Case Study: Repairing Parent Pointers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Online reconstruction of a file's parent pointer information works similarly to
+directory reconstruction:
+
+1. Set up a temporary file for generating a new extended attribute structure,
+   an xfblob for storing parent pointer names, and an xfarray for stashing the
+   fixed size fields involved in a parent pointer update: ``(parent inumber,
+   parent generation, add vs. remove, name cookie)``.
+
+2. Set up an inode scanner and hook into the directory entry code to receive
+   updates on directory operations.
+
+3. For each directory entry found in each directory scanned, decide if the
+   dirent references the file of interest.
+   If so:
+
+   a. Stash the dirent name and an addpptr entry for this parent pointer in the
+      xfblob and xfarray, respectively.
+
+   b. When finished scanning the directory or the kernel memory consumption
+      exceeds a threshold, flush the stashed updates to the temporary file.
+
+4. For each live directory update received via the hook, decide if the parent
+   has already been scanned.
+   If so:
+
+   a. Stash the dirent name and an addpptr or removepptr entry for this dirent
+      update in the xfblob and xfarray for later.
+      We cannot write parent pointers directly to the temporary file because
+      hook functions are not allowed to modify filesystem metadata.
+      Instead, we stash updates in the xfarray and rely on the scanner thread
+      to apply the stashed parent pointer updates to the temporary file.
+
+5. When the scan is complete, replay any stashed entries in the xfarray.
+
+6. Copy all non-parent pointer extended attributes to the temporary file.
+
+7. When the scan is complete, atomically exchange the mappings of the attribute
+   forks of the temporary file and the file being repaired.
+   The temporary file now contains the damaged extended attribute structure.
+
+8. Reap the temporary file.
+
+Digression: Offline Checking of Parent Pointers
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Examining parent pointers in offline repair works differently because corrupt
+files are erased long before directory tree connectivity checks are performed.
+Parent pointer checks are therefore a second pass to be added to the existing
+connectivity checks:
+
+1. After the set of surviving files has been established (phase 6),
+   walk the surviving directories of each AG in the filesystem.
+   This is already performed as part of the connectivity checks.
+
+2. For each directory entry found,
+
+   a. If the name has already been stored in the xfblob, then use that cookie
+      and skip the next step.
+
+   b. Otherwise, record the name in an xfblob, and remember the xfblob cookie.
+      Unique mappings are critical for
+
+      1. Deduplicating names to reduce memory usage, and
+
+      2. Creating a stable sort key for the parent pointer indexes so that the
+         parent pointer validation described below will work.
+
+   c. Store ``(child_ag_inum, parent_inum, parent_gen, name_hash, name_len,
+      name_cookie)`` tuples in a per-AG in-memory slab.  The ``name_hash``
+      referenced in this section is the regular directory entry name hash, not
+      the specialized one used for parent pointer xattrs.
+
+3. For each AG in the filesystem,
+
+   a. Sort the per-AG tuple set in order of ``child_ag_inum``, ``parent_inum``,
+      ``name_hash``, and ``name_cookie``.
+      Having a single ``name_cookie`` for each ``name`` is critical for
+      handling the uncommon case of a directory containing multiple hardlinks
+      to the same file where all the names hash to the same value.
+
+   b. For each inode in the AG,
+
+      1. Scan the inode for parent pointers.
+         For each parent pointer found,
+
+         a. Validate the ondisk parent pointer.
+            If validation fails, move on to the next parent pointer in the
+            file.
+
+         b. If the name has already been stored in the xfblob, then use that
+            cookie and skip the next step.
+
+         c. Record the name in a per-file xfblob, and remember the xfblob
+            cookie.
+
+         d. Store ``(parent_inum, parent_gen, name_hash, name_len,
+            name_cookie)`` tuples in a per-file slab.
+
+      2. Sort the per-file tuples in order of ``parent_inum``, ``name_hash``,
+         and ``name_cookie``.
+
+      3. Position one slab cursor at the start of the inode's records in the
+         per-AG tuple slab.
+         This should be trivial since the per-AG tuples are in child inumber
+         order.
+
+      4. Position a second slab cursor at the start of the per-file tuple slab.
+
+      5. Iterate the two cursors in lockstep, comparing the ``parent_ino``,
+         ``name_hash``, and ``name_cookie`` fields of the records under each
+         cursor:
+
+         a. If the per-AG cursor is at a lower point in the keyspace than the
+            per-file cursor, then the per-AG cursor points to a missing parent
+            pointer.
+            Add the parent pointer to the inode and advance the per-AG
+            cursor.
+
+         b. If the per-file cursor is at a lower point in the keyspace than
+            the per-AG cursor, then the per-file cursor points to a dangling
+            parent pointer.
+            Remove the parent pointer from the inode and advance the per-file
+            cursor.
+
+         c. Otherwise, both cursors point at the same parent pointer.
+            Update the parent_gen component if necessary.
+            Advance both cursors.
+
+4. Move on to examining link counts, as we do today.
+
+Rebuilding directories from parent pointers in offline repair would be very
+challenging because xfs_repair currently uses two single-pass scans of the
+filesystem during phases 3 and 4 to decide which files are corrupt enough to be
+zapped.
+This scan would have to be converted into a multi-pass scan:
+
+1. The first pass of the scan zaps corrupt inodes, forks, and attributes
+   much as it does now.
+   Corrupt directories are noted but not zapped.
+
+2. The next pass records parent pointers pointing to the directories noted
+   as being corrupt in the first pass.
+   This second pass may have to happen after the phase 4 scan for duplicate
+   blocks, if phase 4 is also capable of zapping directories.
+
+3. The third pass resets corrupt directories to an empty shortform directory.
+   Free space metadata has not been ensured yet, so repair cannot yet use the
+   directory building code in libxfs.
+
+4. At the start of phase 6, space metadata have been rebuilt.
+   Use the parent pointer information recorded during step 2 to reconstruct
+   the dirents and add them to the now-empty directories.
+
+This code has not yet been constructed.
+
+.. _dirtree:
+
+Case Study: Directory Tree Structure
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+As mentioned earlier, the filesystem directory tree is supposed to be a
+directed acylic graph structure.
+However, each node in this graph is a separate ``xfs_inode`` object with its
+own locks, which makes validating the tree qualities difficult.
+Fortunately, non-directories are allowed to have multiple parents and cannot
+have children, so only directories need to be scanned.
+Directories typically constitute 5-10% of the files in a filesystem, which
+reduces the amount of work dramatically.
+
+If the directory tree could be frozen, it would be easy to discover cycles and
+disconnected regions by running a depth (or breadth) first search downwards
+from the root directory and marking a bitmap for each directory found.
+At any point in the walk, trying to set an already set bit means there is a
+cycle.
+After the scan completes, XORing the marked inode bitmap with the inode
+allocation bitmap reveals disconnected inodes.
+However, one of online repair's design goals is to avoid locking the entire
+filesystem unless it's absolutely necessary.
+Directory tree updates can move subtrees across the scanner wavefront on a live
+filesystem, so the bitmap algorithm cannot be applied.
+
+Directory parent pointers enable an incremental approach to validation of the
+tree structure.
+Instead of using one thread to scan the entire filesystem, multiple threads can
+walk from individual subdirectories upwards towards the root.
+For this to work, all directory entries and parent pointers must be internally
+consistent, each directory entry must have a parent pointer, and the link
+counts of all directories must be correct.
+Each scanner thread must be able to take the IOLOCK of an alleged parent
+directory while holding the IOLOCK of the child directory to prevent either
+directory from being moved within the tree.
+This is not possible since the VFS does not take the IOLOCK of a child
+subdirectory when moving that subdirectory, so instead the scanner stabilizes
+the parent -> child relationship by taking the ILOCKs and installing a dirent
+update hook to detect changes.
+
+The scanning process uses a dirent hook to detect changes to the directories
+mentioned in the scan data.
+The scan works as follows:
+
+1. For each subdirectory in the filesystem,
+
+   a. For each parent pointer of that subdirectory,
+
+      1. Create a path object for that parent pointer, and mark the
+         subdirectory inode number in the path object's bitmap.
+
+      2. Record the parent pointer name and inode number in a path structure.
+
+      3. If the alleged parent is the subdirectory being scrubbed, the path is
+         a cycle.
+         Mark the path for deletion and repeat step 1a with the next
+         subdirectory parent pointer.
+
+      4. Try to mark the alleged parent inode number in a bitmap in the path
+         object.
+         If the bit is already set, then there is a cycle in the directory
+         tree.
+         Mark the path as a cycle and repeat step 1a with the next subdirectory
+         parent pointer.
+
+      5. Load the alleged parent.
+         If the alleged parent is not a linked directory, abort the scan
+         because the parent pointer information is inconsistent.
+
+      6. For each parent pointer of this alleged ancestor directory,
+
+         a. Record the parent pointer name and inode number in the path object
+            if no parent has been set for that level.
+
+         b. If an ancestor has more than one parent, mark the path as corrupt.
+            Repeat step 1a with the next subdirectory parent pointer.
+
+         c. Repeat steps 1a3-1a6 for the ancestor identified in step 1a6a.
+            This repeats until the directory tree root is reached or no parents
+            are found.
+
+      7. If the walk terminates at the root directory, mark the path as ok.
+
+      8. If the walk terminates without reaching the root, mark the path as
+         disconnected.
+
+2. If the directory entry update hook triggers, check all paths already found
+   by the scan.
+   If the entry matches part of a path, mark that path and the scan stale.
+   When the scanner thread sees that the scan has been marked stale, it deletes
+   all scan data and starts over.
+
+Repairing the directory tree works as follows:
+
+1. Walk each path of the target subdirectory.
+
+   a. Corrupt paths and cycle paths are counted as suspect.
+
+   b. Paths already marked for deletion are counted as bad.
+
+   c. Paths that reached the root are counted as good.
+
+2. If the subdirectory is either the root directory or has zero link count,
+   delete all incoming directory entries in the immediate parents.
+   Repairs are complete.
+
+3. If the subdirectory has exactly one path, set the dotdot entry to the
+   parent and exit.
+
+4. If the subdirectory has at least one good path, delete all the other
+   incoming directory entries in the immediate parents.
+
+5. If the subdirectory has no good paths and more than one suspect path, delete
+   all the other incoming directory entries in the immediate parents.
+
+6. If the subdirectory has zero paths, attach it to the lost and found.
+
+.. _orphanage:
+
+The Orphanage
+-------------
+
+Filesystems present files as a directed, and hopefully acyclic, graph.
+In other words, a tree.
+The root of the filesystem is a directory, and each entry in a directory points
+downwards either to more subdirectories or to non-directory files.
+Unfortunately, a disruption in the directory graph pointers result in a
+disconnected graph, which makes files impossible to access via regular path
+resolution.
+
+Without parent pointers, the directory parent pointer online scrub code can
+detect a dotdot entry pointing to a parent directory that doesn't have a link
+back to the child directory and the file link count checker can detect a file
+that isn't pointed to by any directory in the filesystem.
+If such a file has a positive link count, the file is an orphan.
+
+With parent pointers, directories can be rebuilt by scanning parent pointers
+and parent pointers can be rebuilt by scanning directories.
+This should reduce the incidence of files ending up in ``/lost+found``.
+
+When orphans are found, they should be reconnected to the directory tree.
+Offline fsck solves the problem by creating a directory ``/lost+found`` to
+serve as an orphanage, and linking orphan files into the orphanage by using the
+inumber as the name.
+Reparenting a file to the orphanage does not reset any of its permissions or
+ACLs.
+
+This process is more involved in the kernel than it is in userspace.
+The directory and file link count repair setup functions must use the regular
+VFS mechanisms to create the orphanage directory with all the necessary
+security attributes and dentry cache entries, just like a regular directory
+tree modification.
+
+Orphaned files are adopted by the orphanage as follows:
+
+1. Call ``xrep_orphanage_try_create`` at the start of the scrub setup function
+   to try to ensure that the lost and found directory actually exists.
+   This also attaches the orphanage directory to the scrub context.
+
+2. If the decision is made to reconnect a file, take the IOLOCK of both the
+   orphanage and the file being reattached.
+   The ``xrep_orphanage_iolock_two`` function follows the inode locking
+   strategy discussed earlier.
+
+3. Use ``xrep_adoption_trans_alloc`` to reserve resources to the repair
+   transaction.
+
+4. Call ``xrep_orphanage_compute_name`` to compute the new name in the
+   orphanage.
+
+5. If the adoption is going to happen, call ``xrep_adoption_reparent`` to
+   reparent the orphaned file into the lost and found and invalidate the dentry
+   cache.
+
+6. Call ``xrep_adoption_finish`` to commit any filesystem updates, release the
+   orphanage ILOCK, and clean the scrub transaction.  Call
+   ``xrep_adoption_commit`` to commit the updates and the scrub transaction.
+
+7. If a runtime error happens, call ``xrep_adoption_cancel`` to release all
+   resources.
+
+6. Userspace Algorithms and Data Structures
+===========================================
+
+This section discusses the key algorithms and data structures of the userspace
+program, ``xfs_scrub``, that provide the ability to drive metadata checks and
+repairs in the kernel, verify file data, and look for other potential problems.
+
+.. _scrubcheck:
+
+Checking Metadata
+-----------------
+
+Recall the :ref:`phases of fsck work<scrubphases>` outlined earlier.
+That structure follows naturally from the data dependencies designed into the
+filesystem from its beginnings in 1993.
+In XFS, there are several groups of metadata dependencies:
+
+a. Filesystem summary counts depend on consistency within the inode indices,
+   the allocation group space btrees, and the realtime volume space
+   information.
+
+b. Quota resource counts depend on consistency within the quota file data
+   forks, inode indices, inode records, and the forks of every file on the
+   system.
+
+c. The naming hierarchy depends on consistency within the directory and
+   extended attribute structures.
+   This includes file link counts.
+
+d. Directories, extended attributes, and file data depend on consistency within
+   the file forks that map directory and extended attribute data to physical
+   storage media.
+
+e. The file forks depends on consistency within inode records and the space
+   metadata indices of the allocation groups and the realtime volume.
+   This includes quota and realtime metadata files.
+
+f. Inode records depends on consistency within the inode metadata indices.
+
+g. Realtime space metadata depend on the inode records and data forks of the
+   realtime metadata inodes.
+
+h. The allocation group metadata indices (free space, inodes, reference count,
+   and reverse mapping btrees) depend on consistency within the AG headers and
+   between all the AG metadata btrees.
+
+i. ``xfs_scrub`` depends on the filesystem being mounted and kernel support
+   for online fsck functionality.
+
+Therefore, a metadata dependency graph is a convenient way to schedule checking
+operations in the ``xfs_scrub`` program:
+
+- Phase 1 checks that the provided path maps to an XFS filesystem and detect
+  the kernel's scrubbing abilities, which validates group (i).
+
+- Phase 2 scrubs groups (g) and (h) in parallel using a threaded workqueue.
+
+- Phase 3 scans inodes in parallel.
+  For each inode, groups (f), (e), and (d) are checked, in that order.
+
+- Phase 4 repairs everything in groups (i) through (d) so that phases 5 and 6
+  may run reliably.
+
+- Phase 5 starts by checking groups (b) and (c) in parallel before moving on
+  to checking names.
+
+- Phase 6 depends on groups (i) through (b) to find file data blocks to verify,
+  to read them, and to report which blocks of which files are affected.
+
+- Phase 7 checks group (a), having validated everything else.
+
+Notice that the data dependencies between groups are enforced by the structure
+of the program flow.
+
+Parallel Inode Scans
+--------------------
+
+An XFS filesystem can easily contain hundreds of millions of inodes.
+Given that XFS targets installations with large high-performance storage,
+it is desirable to scrub inodes in parallel to minimize runtime, particularly
+if the program has been invoked manually from a command line.
+This requires careful scheduling to keep the threads as evenly loaded as
+possible.
+
+Early iterations of the ``xfs_scrub`` inode scanner naïvely created a single
+workqueue and scheduled a single workqueue item per AG.
+Each workqueue item walked the inode btree (with ``XFS_IOC_INUMBERS``) to find
+inode chunks and then called bulkstat (``XFS_IOC_BULKSTAT``) to gather enough
+information to construct file handles.
+The file handle was then passed to a function to generate scrub items for each
+metadata object of each inode.
+This simple algorithm leads to thread balancing problems in phase 3 if the
+filesystem contains one AG with a few large sparse files and the rest of the
+AGs contain many smaller files.
+The inode scan dispatch function was not sufficiently granular; it should have
+been dispatching at the level of individual inodes, or, to constrain memory
+consumption, inode btree records.
+
+Thanks to Dave Chinner, bounded workqueues in userspace enable ``xfs_scrub`` to
+avoid this problem with ease by adding a second workqueue.
+Just like before, the first workqueue is seeded with one workqueue item per AG,
+and it uses INUMBERS to find inode btree chunks.
+The second workqueue, however, is configured with an upper bound on the number
+of items that can be waiting to be run.
+Each inode btree chunk found by the first workqueue's workers are queued to the
+second workqueue, and it is this second workqueue that queries BULKSTAT,
+creates a file handle, and passes it to a function to generate scrub items for
+each metadata object of each inode.
+If the second workqueue is too full, the workqueue add function blocks the
+first workqueue's workers until the backlog eases.
+This doesn't completely solve the balancing problem, but reduces it enough to
+move on to more pressing issues.
+
+.. _scrubrepair:
+
+Scheduling Repairs
+------------------
+
+During phase 2, corruptions and inconsistencies reported in any AGI header or
+inode btree are repaired immediately, because phase 3 relies on proper
+functioning of the inode indices to find inodes to scan.
+Failed repairs are rescheduled to phase 4.
+Problems reported in any other space metadata are deferred to phase 4.
+Optimization opportunities are always deferred to phase 4, no matter their
+origin.
+
+During phase 3, corruptions and inconsistencies reported in any part of a
+file's metadata are repaired immediately if all space metadata were validated
+during phase 2.
+Repairs that fail or cannot be repaired immediately are scheduled for phase 4.
+
+In the original design of ``xfs_scrub``, it was thought that repairs would be
+so infrequent that the ``struct xfs_scrub_metadata`` objects used to
+communicate with the kernel could also be used as the primary object to
+schedule repairs.
+With recent increases in the number of optimizations possible for a given
+filesystem object, it became much more memory-efficient to track all eligible
+repairs for a given filesystem object with a single repair item.
+Each repair item represents a single lockable object -- AGs, metadata files,
+individual inodes, or a class of summary information.
+
+Phase 4 is responsible for scheduling a lot of repair work in as quick a
+manner as is practical.
+The :ref:`data dependencies <scrubcheck>` outlined earlier still apply, which
+means that ``xfs_scrub`` must try to complete the repair work scheduled by
+phase 2 before trying repair work scheduled by phase 3.
+The repair process is as follows:
+
+1. Start a round of repair with a workqueue and enough workers to keep the CPUs
+   as busy as the user desires.
+
+   a. For each repair item queued by phase 2,
+
+      i.   Ask the kernel to repair everything listed in the repair item for a
+           given filesystem object.
+
+      ii.  Make a note if the kernel made any progress in reducing the number
+           of repairs needed for this object.
+
+      iii. If the object no longer requires repairs, revalidate all metadata
+           associated with this object.
+           If the revalidation succeeds, drop the repair item.
+           If not, requeue the item for more repairs.
+
+   b. If any repairs were made, jump back to 1a to retry all the phase 2 items.
+
+   c. For each repair item queued by phase 3,
+
+      i.   Ask the kernel to repair everything listed in the repair item for a
+           given filesystem object.
+
+      ii.  Make a note if the kernel made any progress in reducing the number
+           of repairs needed for this object.
+
+      iii. If the object no longer requires repairs, revalidate all metadata
+           associated with this object.
+           If the revalidation succeeds, drop the repair item.
+           If not, requeue the item for more repairs.
+
+   d. If any repairs were made, jump back to 1c to retry all the phase 3 items.
+
+2. If step 1 made any repair progress of any kind, jump back to step 1 to start
+   another round of repair.
+
+3. If there are items left to repair, run them all serially one more time.
+   Complain if the repairs were not successful, since this is the last chance
+   to repair anything.
+
+Corruptions and inconsistencies encountered during phases 5 and 7 are repaired
+immediately.
+Corrupt file data blocks reported by phase 6 cannot be recovered by the
+filesystem.
+
+Checking Names for Confusable Unicode Sequences
+-----------------------------------------------
+
+If ``xfs_scrub`` succeeds in validating the filesystem metadata by the end of
+phase 4, it moves on to phase 5, which checks for suspicious looking names in
+the filesystem.
+These names consist of the filesystem label, names in directory entries, and
+the names of extended attributes.
+Like most Unix filesystems, XFS imposes the sparest of constraints on the
+contents of a name:
+
+- Slashes and null bytes are not allowed in directory entries.
+
+- Null bytes are not allowed in userspace-visible extended attributes.
+
+- Null bytes are not allowed in the filesystem label.
+
+Directory entries and attribute keys store the length of the name explicitly
+ondisk, which means that nulls are not name terminators.
+For this section, the term "naming domain" refers to any place where names are
+presented together -- all the names in a directory, or all the attributes of a
+file.
+
+Although the Unix naming constraints are very permissive, the reality of most
+modern-day Linux systems is that programs work with Unicode character code
+points to support international languages.
+These programs typically encode those code points in UTF-8 when interfacing
+with the C library because the kernel expects null-terminated names.
+In the common case, therefore, names found in an XFS filesystem are actually
+UTF-8 encoded Unicode data.
+
+To maximize its expressiveness, the Unicode standard defines separate control
+points for various characters that render similarly or identically in writing
+systems around the world.
+For example, the character "Cyrillic Small Letter A" U+0430 "а" often renders
+identically to "Latin Small Letter A" U+0061 "a".
+
+The standard also permits characters to be constructed in multiple ways --
+either by using a defined code point, or by combining one code point with
+various combining marks.
+For example, the character "Angstrom Sign U+212B "Å" can also be expressed
+as "Latin Capital Letter A" U+0041 "A" followed by "Combining Ring Above"
+U+030A "◌̊".
+Both sequences render identically.
+
+Like the standards that preceded it, Unicode also defines various control
+characters to alter the presentation of text.
+For example, the character "Right-to-Left Override" U+202E can trick some
+programs into rendering "moo\\xe2\\x80\\xaegnp.txt" as "mootxt.png".
+A second category of rendering problems involves whitespace characters.
+If the character "Zero Width Space" U+200B is encountered in a file name, the
+name will render identically to a name that does not have the zero width
+space.
+
+If two names within a naming domain have different byte sequences but render
+identically, a user may be confused by it.
+The kernel, in its indifference to upper level encoding schemes, permits this.
+Most filesystem drivers persist the byte sequence names that are given to them
+by the VFS.
+
+Techniques for detecting confusable names are explained in great detail in
+sections 4 and 5 of the
+`Unicode Security Mechanisms <https://unicode.org/reports/tr39/>`_
+document.
+When ``xfs_scrub`` detects UTF-8 encoding in use on a system, it uses the
+Unicode normalization form NFD in conjunction with the confusable name
+detection component of
+`libicu <https://github.com/unicode-org/icu>`_
+to identify names with a directory or within a file's extended attributes that
+could be confused for each other.
+Names are also checked for control characters, non-rendering characters, and
+mixing of bidirectional characters.
+All of these potential issues are reported to the system administrator during
+phase 5.
+
+Media Verification of File Data Extents
+---------------------------------------
+
+The system administrator can elect to initiate a media scan of all file data
+blocks.
+This scan after validation of all filesystem metadata (except for the summary
+counters) as phase 6.
+The scan starts by calling ``FS_IOC_GETFSMAP`` to scan the filesystem space map
+to find areas that are allocated to file data fork extents.
+Gaps between data fork extents that are smaller than 64k are treated as if
+they were data fork extents to reduce the command setup overhead.
+When the space map scan accumulates a region larger than 32MB, a media
+verification request is sent to the disk as a directio read of the raw block
+device.
+
+If the verification read fails, ``xfs_scrub`` retries with single-block reads
+to narrow down the failure to the specific region of the media and recorded.
+When it has finished issuing verification requests, it again uses the space
+mapping ioctl to map the recorded media errors back to metadata structures
+and report what has been lost.
+For media errors in blocks owned by files, parent pointers can be used to
+construct file paths from inode numbers for user-friendly reporting.
+
+7. Conclusion and Future Work
+=============================
+
+It is hoped that the reader of this document has followed the designs laid out
+in this document and now has some familiarity with how XFS performs online
+rebuilding of its metadata indices, and how filesystem users can interact with
+that functionality.
+Although the scope of this work is daunting, it is hoped that this guide will
+make it easier for code readers to understand what has been built, for whom it
+has been built, and why.
+Please feel free to contact the XFS mailing list with questions.
+
+XFS_IOC_EXCHANGE_RANGE
+----------------------
+
+As discussed earlier, a second frontend to the atomic file mapping exchange
+mechanism is a new ioctl call that userspace programs can use to commit updates
+to files atomically.
+This frontend has been out for review for several years now, though the
+necessary refinements to online repair and lack of customer demand mean that
+the proposal has not been pushed very hard.
+
+File Content Exchanges with Regular User Files
+``````````````````````````````````````````````
+
+As mentioned earlier, XFS has long had the ability to swap extents between
+files, which is used almost exclusively by ``xfs_fsr`` to defragment files.
+The earliest form of this was the fork swap mechanism, where the entire
+contents of data forks could be exchanged between two files by exchanging the
+raw bytes in each inode fork's immediate area.
+When XFS v5 came along with self-describing metadata, this old mechanism grew
+some log support to continue rewriting the owner fields of BMBT blocks during
+log recovery.
+When the reverse mapping btree was later added to XFS, the only way to maintain
+the consistency of the fork mappings with the reverse mapping index was to
+develop an iterative mechanism that used deferred bmap and rmap operations to
+swap mappings one at a time.
+This mechanism is identical to steps 2-3 from the procedure above except for
+the new tracking items, because the atomic file mapping exchange mechanism is
+an iteration of an existing mechanism and not something totally novel.
+For the narrow case of file defragmentation, the file contents must be
+identical, so the recovery guarantees are not much of a gain.
+
+Atomic file content exchanges are much more flexible than the existing swapext
+implementations because it can guarantee that the caller never sees a mix of
+old and new contents even after a crash, and it can operate on two arbitrary
+file fork ranges.
+The extra flexibility enables several new use cases:
+
+- **Atomic commit of file writes**: A userspace process opens a file that it
+  wants to update.
+  Next, it opens a temporary file and calls the file clone operation to reflink
+  the first file's contents into the temporary file.
+  Writes to the original file should instead be written to the temporary file.
+  Finally, the process calls the atomic file mapping exchange system call
+  (``XFS_IOC_EXCHANGE_RANGE``) to exchange the file contents, thereby
+  committing all of the updates to the original file, or none of them.
+
+.. _exchrange_if_unchanged:
+
+- **Transactional file updates**: The same mechanism as above, but the caller
+  only wants the commit to occur if the original file's contents have not
+  changed.
+  To make this happen, the calling process snapshots the file modification and
+  change timestamps of the original file before reflinking its data to the
+  temporary file.
+  When the program is ready to commit the changes, it passes the timestamps
+  into the kernel as arguments to the atomic file mapping exchange system call.
+  The kernel only commits the changes if the provided timestamps match the
+  original file.
+  A new ioctl (``XFS_IOC_COMMIT_RANGE``) is provided to perform this.
+
+- **Emulation of atomic block device writes**: Export a block device with a
+  logical sector size matching the filesystem block size to force all writes
+  to be aligned to the filesystem block size.
+  Stage all writes to a temporary file, and when that is complete, call the
+  atomic file mapping exchange system call with a flag to indicate that holes
+  in the temporary file should be ignored.
+  This emulates an atomic device write in software, and can support arbitrary
+  scattered writes.
+
+(This functionality was merged into mainline as of 2025)
+
+Vectorized Scrub
+----------------
+
+As it turns out, the :ref:`refactoring <scrubrepair>` of repair items mentioned
+earlier was a catalyst for enabling a vectorized scrub system call.
+Since 2018, the cost of making a kernel call has increased considerably on some
+systems to mitigate the effects of speculative execution attacks.
+This incentivizes program authors to make as few system calls as possible to
+reduce the number of times an execution path crosses a security boundary.
+
+With vectorized scrub, userspace pushes to the kernel the identity of a
+filesystem object, a list of scrub types to run against that object, and a
+simple representation of the data dependencies between the selected scrub
+types.
+The kernel executes as much of the caller's plan as it can until it hits a
+dependency that cannot be satisfied due to a corruption, and tells userspace
+how much was accomplished.
+It is hoped that ``io_uring`` will pick up enough of this functionality that
+online fsck can use that instead of adding a separate vectored scrub system
+call to XFS.
+
+(This functionality was merged into mainline as of 2025)
+
+Quality of Service Targets for Scrub
+------------------------------------
+
+One serious shortcoming of the online fsck code is that the amount of time that
+it can spend in the kernel holding resource locks is basically unbounded.
+Userspace is allowed to send a fatal signal to the process which will cause
+``xfs_scrub`` to exit when it reaches a good stopping point, but there's no way
+for userspace to provide a time budget to the kernel.
+Given that the scrub codebase has helpers to detect fatal signals, it shouldn't
+be too much work to allow userspace to specify a timeout for a scrub/repair
+operation and abort the operation if it exceeds budget.
+However, most repair functions have the property that once they begin to touch
+ondisk metadata, the operation cannot be cancelled cleanly, after which a QoS
+timeout is no longer useful.
+
+Defragmenting Free Space
+------------------------
+
+Over the years, many XFS users have requested the creation of a program to
+clear a portion of the physical storage underlying a filesystem so that it
+becomes a contiguous chunk of free space.
+Call this free space defragmenter ``clearspace`` for short.
+
+The first piece the ``clearspace`` program needs is the ability to read the
+reverse mapping index from userspace.
+This already exists in the form of the ``FS_IOC_GETFSMAP`` ioctl.
+The second piece it needs is a new fallocate mode
+(``FALLOC_FL_MAP_FREE_SPACE``) that allocates the free space in a region and
+maps it to a file.
+Call this file the "space collector" file.
+The third piece is the ability to force an online repair.
+
+To clear all the metadata out of a portion of physical storage, clearspace
+uses the new fallocate map-freespace call to map any free space in that region
+to the space collector file.
+Next, clearspace finds all metadata blocks in that region by way of
+``GETFSMAP`` and issues forced repair requests on the data structure.
+This often results in the metadata being rebuilt somewhere that is not being
+cleared.
+After each relocation, clearspace calls the "map free space" function again to
+collect any newly freed space in the region being cleared.
+
+To clear all the file data out of a portion of the physical storage, clearspace
+uses the FSMAP information to find relevant file data blocks.
+Having identified a good target, it uses the ``FICLONERANGE`` call on that part
+of the file to try to share the physical space with a dummy file.
+Cloning the extent means that the original owners cannot overwrite the
+contents; any changes will be written somewhere else via copy-on-write.
+Clearspace makes its own copy of the frozen extent in an area that is not being
+cleared, and uses ``FIEDEUPRANGE`` (or the :ref:`atomic file content exchanges
+<exchrange_if_unchanged>` feature) to change the target file's data extent
+mapping away from the area being cleared.
+When all other mappings have been moved, clearspace reflinks the space into the
+space collector file so that it becomes unavailable.
+
+There are further optimizations that could apply to the above algorithm.
+To clear a piece of physical storage that has a high sharing factor, it is
+strongly desirable to retain this sharing factor.
+In fact, these extents should be moved first to maximize sharing factor after
+the operation completes.
+To make this work smoothly, clearspace needs a new ioctl
+(``FS_IOC_GETREFCOUNTS``) to report reference count information to userspace.
+With the refcount information exposed, clearspace can quickly find the longest,
+most shared data extents in the filesystem, and target them first.
+
+**Future Work Question**: How might the filesystem move inode chunks?
+
+*Answer*: To move inode chunks, Dave Chinner constructed a prototype program
+that creates a new file with the old contents and then locklessly runs around
+the filesystem updating directory entries.
+The operation cannot complete if the filesystem goes down.
+That problem isn't totally insurmountable: create an inode remapping table
+hidden behind a jump label, and a log item that tracks the kernel walking the
+filesystem to update directory entries.
+The trouble is, the kernel can't do anything about open files, since it cannot
+revoke them.
+
+**Future Work Question**: Can static keys be used to minimize the cost of
+supporting ``revoke()`` on XFS files?
+
+*Answer*: Yes.
+Until the first revocation, the bailout code need not be in the call path at
+all.
+
+The relevant patchsets are the
+`kernel freespace defrag
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=defrag-freespace>`_
+and
+`userspace freespace defrag
+<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=defrag-freespace>`_
+series.
+
+Shrinking Filesystems
+---------------------
+
+Removing the end of the filesystem ought to be a simple matter of evacuating
+the data and metadata at the end of the filesystem, and handing the freed space
+to the shrink code.
+That requires an evacuation of the space at end of the filesystem, which is a
+use of free space defragmentation!
diff --git a/Documentation/filesystems/xfs/xfs-self-describing-metadata.rst b/Documentation/filesystems/xfs/xfs-self-describing-metadata.rst
new file mode 100644
index 000000000..a10c4ae69
--- /dev/null
+++ b/Documentation/filesystems/xfs/xfs-self-describing-metadata.rst
@@ -0,0 +1,353 @@
+.. SPDX-License-Identifier: GPL-2.0
+.. _xfs_self_describing_metadata:
+
+============================
+XFS Self Describing Metadata
+============================
+
+Introduction
+============
+
+The largest scalability problem facing XFS is not one of algorithmic
+scalability, but of verification of the filesystem structure. Scalabilty of the
+structures and indexes on disk and the algorithms for iterating them are
+adequate for supporting PB scale filesystems with billions of inodes, however it
+is this very scalability that causes the verification problem.
+
+Almost all metadata on XFS is dynamically allocated. The only fixed location
+metadata is the allocation group headers (SB, AGF, AGFL and AGI), while all
+other metadata structures need to be discovered by walking the filesystem
+structure in different ways. While this is already done by userspace tools for
+validating and repairing the structure, there are limits to what they can
+verify, and this in turn limits the supportable size of an XFS filesystem.
+
+For example, it is entirely possible to manually use xfs_db and a bit of
+scripting to analyse the structure of a 100TB filesystem when trying to
+determine the root cause of a corruption problem, but it is still mainly a
+manual task of verifying that things like single bit errors or misplaced writes
+weren't the ultimate cause of a corruption event. It may take a few hours to a
+few days to perform such forensic analysis, so for at this scale root cause
+analysis is entirely possible.
+
+However, if we scale the filesystem up to 1PB, we now have 10x as much metadata
+to analyse and so that analysis blows out towards weeks/months of forensic work.
+Most of the analysis work is slow and tedious, so as the amount of analysis goes
+up, the more likely that the cause will be lost in the noise.  Hence the primary
+concern for supporting PB scale filesystems is minimising the time and effort
+required for basic forensic analysis of the filesystem structure.
+
+
+Self Describing Metadata
+========================
+
+One of the problems with the current metadata format is that apart from the
+magic number in the metadata block, we have no other way of identifying what it
+is supposed to be. We can't even identify if it is the right place. Put simply,
+you can't look at a single metadata block in isolation and say "yes, it is
+supposed to be there and the contents are valid".
+
+Hence most of the time spent on forensic analysis is spent doing basic
+verification of metadata values, looking for values that are in range (and hence
+not detected by automated verification checks) but are not correct. Finding and
+understanding how things like cross linked block lists (e.g. sibling
+pointers in a btree end up with loops in them) are the key to understanding what
+went wrong, but it is impossible to tell what order the blocks were linked into
+each other or written to disk after the fact.
+
+Hence we need to record more information into the metadata to allow us to
+quickly determine if the metadata is intact and can be ignored for the purpose
+of analysis. We can't protect against every possible type of error, but we can
+ensure that common types of errors are easily detectable.  Hence the concept of
+self describing metadata.
+
+The first, fundamental requirement of self describing metadata is that the
+metadata object contains some form of unique identifier in a well known
+location. This allows us to identify the expected contents of the block and
+hence parse and verify the metadata object. IF we can't independently identify
+the type of metadata in the object, then the metadata doesn't describe itself
+very well at all!
+
+Luckily, almost all XFS metadata has magic numbers embedded already - only the
+AGFL, remote symlinks and remote attribute blocks do not contain identifying
+magic numbers. Hence we can change the on-disk format of all these objects to
+add more identifying information and detect this simply by changing the magic
+numbers in the metadata objects. That is, if it has the current magic number,
+the metadata isn't self identifying. If it contains a new magic number, it is
+self identifying and we can do much more expansive automated verification of the
+metadata object at runtime, during forensic analysis or repair.
+
+As a primary concern, self describing metadata needs some form of overall
+integrity checking. We cannot trust the metadata if we cannot verify that it has
+not been changed as a result of external influences. Hence we need some form of
+integrity check, and this is done by adding CRC32c validation to the metadata
+block. If we can verify the block contains the metadata it was intended to
+contain, a large amount of the manual verification work can be skipped.
+
+CRC32c was selected as metadata cannot be more than 64k in length in XFS and
+hence a 32 bit CRC is more than sufficient to detect multi-bit errors in
+metadata blocks. CRC32c is also now hardware accelerated on common CPUs so it is
+fast. So while CRC32c is not the strongest of possible integrity checks that
+could be used, it is more than sufficient for our needs and has relatively
+little overhead. Adding support for larger integrity fields and/or algorithms
+does really provide any extra value over CRC32c, but it does add a lot of
+complexity and so there is no provision for changing the integrity checking
+mechanism.
+
+Self describing metadata needs to contain enough information so that the
+metadata block can be verified as being in the correct place without needing to
+look at any other metadata. This means it needs to contain location information.
+Just adding a block number to the metadata is not sufficient to protect against
+mis-directed writes - a write might be misdirected to the wrong LUN and so be
+written to the "correct block" of the wrong filesystem. Hence location
+information must contain a filesystem identifier as well as a block number.
+
+Another key information point in forensic analysis is knowing who the metadata
+block belongs to. We already know the type, the location, that it is valid
+and/or corrupted, and how long ago that it was last modified. Knowing the owner
+of the block is important as it allows us to find other related metadata to
+determine the scope of the corruption. For example, if we have a extent btree
+object, we don't know what inode it belongs to and hence have to walk the entire
+filesystem to find the owner of the block. Worse, the corruption could mean that
+no owner can be found (i.e. it's an orphan block), and so without an owner field
+in the metadata we have no idea of the scope of the corruption. If we have an
+owner field in the metadata object, we can immediately do top down validation to
+determine the scope of the problem.
+
+Different types of metadata have different owner identifiers. For example,
+directory, attribute and extent tree blocks are all owned by an inode, while
+freespace btree blocks are owned by an allocation group. Hence the size and
+contents of the owner field are determined by the type of metadata object we are
+looking at.  The owner information can also identify misplaced writes (e.g.
+freespace btree block written to the wrong AG).
+
+Self describing metadata also needs to contain some indication of when it was
+written to the filesystem. One of the key information points when doing forensic
+analysis is how recently the block was modified. Correlation of set of corrupted
+metadata blocks based on modification times is important as it can indicate
+whether the corruptions are related, whether there's been multiple corruption
+events that lead to the eventual failure, and even whether there are corruptions
+present that the run-time verification is not detecting.
+
+For example, we can determine whether a metadata object is supposed to be free
+space or still allocated if it is still referenced by its owner by looking at
+when the free space btree block that contains the block was last written
+compared to when the metadata object itself was last written.  If the free space
+block is more recent than the object and the object's owner, then there is a
+very good chance that the block should have been removed from the owner.
+
+To provide this "written timestamp", each metadata block gets the Log Sequence
+Number (LSN) of the most recent transaction it was modified on written into it.
+This number will always increase over the life of the filesystem, and the only
+thing that resets it is running xfs_repair on the filesystem. Further, by use of
+the LSN we can tell if the corrupted metadata all belonged to the same log
+checkpoint and hence have some idea of how much modification occurred between
+the first and last instance of corrupt metadata on disk and, further, how much
+modification occurred between the corruption being written and when it was
+detected.
+
+Runtime Validation
+==================
+
+Validation of self-describing metadata takes place at runtime in two places:
+
+	- immediately after a successful read from disk
+	- immediately prior to write IO submission
+
+The verification is completely stateless - it is done independently of the
+modification process, and seeks only to check that the metadata is what it says
+it is and that the metadata fields are within bounds and internally consistent.
+As such, we cannot catch all types of corruption that can occur within a block
+as there may be certain limitations that operational state enforces of the
+metadata, or there may be corruption of interblock relationships (e.g. corrupted
+sibling pointer lists). Hence we still need stateful checking in the main code
+body, but in general most of the per-field validation is handled by the
+verifiers.
+
+For read verification, the caller needs to specify the expected type of metadata
+that it should see, and the IO completion process verifies that the metadata
+object matches what was expected. If the verification process fails, then it
+marks the object being read as EFSCORRUPTED. The caller needs to catch this
+error (same as for IO errors), and if it needs to take special action due to a
+verification error it can do so by catching the EFSCORRUPTED error value. If we
+need more discrimination of error type at higher levels, we can define new
+error numbers for different errors as necessary.
+
+The first step in read verification is checking the magic number and determining
+whether CRC validating is necessary. If it is, the CRC32c is calculated and
+compared against the value stored in the object itself. Once this is validated,
+further checks are made against the location information, followed by extensive
+object specific metadata validation. If any of these checks fail, then the
+buffer is considered corrupt and the EFSCORRUPTED error is set appropriately.
+
+Write verification is the opposite of the read verification - first the object
+is extensively verified and if it is OK we then update the LSN from the last
+modification made to the object, After this, we calculate the CRC and insert it
+into the object. Once this is done the write IO is allowed to continue. If any
+error occurs during this process, the buffer is again marked with a EFSCORRUPTED
+error for the higher layers to catch.
+
+Structures
+==========
+
+A typical on-disk structure needs to contain the following information::
+
+    struct xfs_ondisk_hdr {
+	    __be32  magic;		/* magic number */
+	    __be32  crc;		/* CRC, not logged */
+	    uuid_t  uuid;		/* filesystem identifier */
+	    __be64  owner;		/* parent object */
+	    __be64  blkno;		/* location on disk */
+	    __be64  lsn;		/* last modification in log, not logged */
+    };
+
+Depending on the metadata, this information may be part of a header structure
+separate to the metadata contents, or may be distributed through an existing
+structure. The latter occurs with metadata that already contains some of this
+information, such as the superblock and AG headers.
+
+Other metadata may have different formats for the information, but the same
+level of information is generally provided. For example:
+
+	- short btree blocks have a 32 bit owner (ag number) and a 32 bit block
+	  number for location. The two of these combined provide the same
+	  information as @owner and @blkno in eh above structure, but using 8
+	  bytes less space on disk.
+
+	- directory/attribute node blocks have a 16 bit magic number, and the
+	  header that contains the magic number has other information in it as
+	  well. hence the additional metadata headers change the overall format
+	  of the metadata.
+
+A typical buffer read verifier is structured as follows::
+
+    #define XFS_FOO_CRC_OFF		offsetof(struct xfs_ondisk_hdr, crc)
+
+    static void
+    xfs_foo_read_verify(
+	    struct xfs_buf	*bp)
+    {
+	struct xfs_mount *mp = bp->b_mount;
+
+	    if ((xfs_sb_version_hascrc(&mp->m_sb) &&
+		!xfs_verify_cksum(bp->b_addr, BBTOB(bp->b_length),
+					    XFS_FOO_CRC_OFF)) ||
+		!xfs_foo_verify(bp)) {
+		    XFS_CORRUPTION_ERROR(__func__, XFS_ERRLEVEL_LOW, mp, bp->b_addr);
+		    xfs_buf_ioerror(bp, EFSCORRUPTED);
+	    }
+    }
+
+The code ensures that the CRC is only checked if the filesystem has CRCs enabled
+by checking the superblock of the feature bit, and then if the CRC verifies OK
+(or is not needed) it verifies the actual contents of the block.
+
+The verifier function will take a couple of different forms, depending on
+whether the magic number can be used to determine the format of the block. In
+the case it can't, the code is structured as follows::
+
+    static bool
+    xfs_foo_verify(
+	    struct xfs_buf		*bp)
+    {
+	    struct xfs_mount	*mp = bp->b_mount;
+	    struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+
+	    if (hdr->magic != cpu_to_be32(XFS_FOO_MAGIC))
+		    return false;
+
+	    if (!xfs_sb_version_hascrc(&mp->m_sb)) {
+		    if (!uuid_equal(&hdr->uuid, &mp->m_sb.sb_uuid))
+			    return false;
+		    if (bp->b_bn != be64_to_cpu(hdr->blkno))
+			    return false;
+		    if (hdr->owner == 0)
+			    return false;
+	    }
+
+	    /* object specific verification checks here */
+
+	    return true;
+    }
+
+If there are different magic numbers for the different formats, the verifier
+will look like::
+
+    static bool
+    xfs_foo_verify(
+	    struct xfs_buf		*bp)
+    {
+	    struct xfs_mount	*mp = bp->b_mount;
+	    struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+
+	    if (hdr->magic == cpu_to_be32(XFS_FOO_CRC_MAGIC)) {
+		    if (!uuid_equal(&hdr->uuid, &mp->m_sb.sb_uuid))
+			    return false;
+		    if (bp->b_bn != be64_to_cpu(hdr->blkno))
+			    return false;
+		    if (hdr->owner == 0)
+			    return false;
+	    } else if (hdr->magic != cpu_to_be32(XFS_FOO_MAGIC))
+		    return false;
+
+	    /* object specific verification checks here */
+
+	    return true;
+    }
+
+Write verifiers are very similar to the read verifiers, they just do things in
+the opposite order to the read verifiers. A typical write verifier::
+
+    static void
+    xfs_foo_write_verify(
+	    struct xfs_buf	*bp)
+    {
+	    struct xfs_mount	*mp = bp->b_mount;
+	    struct xfs_buf_log_item	*bip = bp->b_fspriv;
+
+	    if (!xfs_foo_verify(bp)) {
+		    XFS_CORRUPTION_ERROR(__func__, XFS_ERRLEVEL_LOW, mp, bp->b_addr);
+		    xfs_buf_ioerror(bp, EFSCORRUPTED);
+		    return;
+	    }
+
+	    if (!xfs_sb_version_hascrc(&mp->m_sb))
+		    return;
+
+
+	    if (bip) {
+		    struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+		    hdr->lsn = cpu_to_be64(bip->bli_item.li_lsn);
+	    }
+	    xfs_update_cksum(bp->b_addr, BBTOB(bp->b_length), XFS_FOO_CRC_OFF);
+    }
+
+This will verify the internal structure of the metadata before we go any
+further, detecting corruptions that have occurred as the metadata has been
+modified in memory. If the metadata verifies OK, and CRCs are enabled, we then
+update the LSN field (when it was last modified) and calculate the CRC on the
+metadata. Once this is done, we can issue the IO.
+
+Inodes and Dquots
+=================
+
+Inodes and dquots are special snowflakes. They have per-object CRC and
+self-identifiers, but they are packed so that there are multiple objects per
+buffer. Hence we do not use per-buffer verifiers to do the work of per-object
+verification and CRC calculations. The per-buffer verifiers simply perform basic
+identification of the buffer - that they contain inodes or dquots, and that
+there are magic numbers in all the expected spots. All further CRC and
+verification checks are done when each inode is read from or written back to the
+buffer.
+
+The structure of the verifiers and the identifiers checks is very similar to the
+buffer code described above. The only difference is where they are called. For
+example, inode read verification is done in xfs_inode_from_disk() when the inode
+is first read out of the buffer and the struct xfs_inode is instantiated. The
+inode is already extensively verified during writeback in xfs_iflush_int, so the
+only addition here is to add the LSN and CRC to the inode as it is copied back
+into the buffer.
+
+XXX: inode unlinked list modification doesn't recalculate the inode CRC! None of
+the unlinked list modifications check or update CRCs, neither during unlink nor
+log recovery. So, it's gone unnoticed until now. This won't matter immediately -
+repair will probably complain about it - but it needs to be fixed.