[PATCH 04/22] docs: add XFS delayed logging design doc to DS&A book

Darrick J. Wong
2018-10-04T04:19:11Z

.../xfs-data-structures/delayed_logging.rst | 828 ++++++++++++++++++++
.../filesystems/xfs-data-structures/overview.rst | 1
.../filesystems/xfs-delayed-logging-design.txt | 793 -------------------
3 files changed, 829 insertions(+), 793 deletions(-)
create mode 100644 Documentation/filesystems/xfs-data-structures/delayed_logging.rst
delete mode 100644 Documentation/filesystems/xfs-delayed-logging-design.txt

diff --git a/Documentation/filesystems/xfs-data-structures/delayed_logging.rst b/Documentation/filesystems/xfs-data-structures/delayed_logging.rst
new file mode 100644
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@@ -0,0 +1,828 @@
+.. SPDX-License-Identifier: CC-BY-SA-4.0
+
+Delayed Logging
+---------------
+
+Introduction to Re-logging in XFS
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+XFS logging is a combination of logical and physical logging. 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. The reason for these differences is to reduce the amount of log space
+required for objects that are frequently logged. Some parts of inodes are more
+frequently logged than others, and inodes are typically more frequently logged
+than any other object (except maybe the superblock buffer) so keeping the
+amount of metadata logged low is of prime importance.
+
+The reason that this is such a concern is that 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
+
+ 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 also 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 - the LSN is effectively a
+direct encoding of the location in the log of the transaction.
+
+This relogging is also used to implement long-running, multiple-commit
+transactions. These transaction are known as rolling transactions, and require
+a special log reservation known as a permanent transaction reservation. 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.
+
+Another feature of the XFS transaction subsystem is that most transactions are
+asynchronous. That is, they don’t commit to disk 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 it’s 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 it’s 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.
+
+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 allow 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.
+
+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.
+
+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.
+
+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 it’s context.
+
+Just to make matters more 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.
+
+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-data-structures/overview.rst b/Documentation/filesystems/xfs-data-structures/overview.rst
index 8b3de9abcf39..457e81c0eb40 100644
--- a/Documentation/filesystems/xfs-data-structures/overview.rst
+++ b/Documentation/filesystems/xfs-data-structures/overview.rst
@@ -44,3 +44,4 @@ are tracked more simply and in larger chunks to reduce jitter in allocation
latency.

.. include:: self_describing_metadata.rst
+.. include:: delayed_logging.rst
diff --git a/Documentation/filesystems/xfs-delayed-logging-design.txt b/Documentation/filesystems/xfs-delayed-logging-design.txt
deleted file mode 100644
index 2ce36439c09f..000000000000
--- a/Documentation/filesystems/xfs-delayed-logging-design.txt
+++ /dev/null
@@ -1,793 +0,0 @@
-XFS Delayed Logging Design
---------------------------
-
-Introduction to Re-logging in XFS
----------------------------------
-
-XFS logging is a combination of logical and physical logging. 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. The reason for these differences is to reduce the amount of log space
-required for objects that are frequently logged. Some parts of inodes are more
-frequently logged than others, and inodes are typically more frequently logged
-than any other object (except maybe the superblock buffer) so keeping the
-amount of metadata logged low is of prime importance.
-
-The reason that this is such a concern is that 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
-
- E E Y (> X+n+m+o)
- F E+F Yٍ+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 also 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 - the LSN is effectively a
-direct encoding of the location in the log of the transaction.
-
-This relogging is also used to implement long-running, multiple-commit
-transactions. These transaction are known as rolling transactions, and require
-a special log reservation known as a permanent transaction reservation. 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.
-
-Another feature of the XFS transaction subsystem is that most transactions are
-asynchronous. That is, they don't commit to disk 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 it's 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
-it's 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 allow 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 it's context.
-
-Just to make matters more 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.