xfs: document future directions of online fsck
Add the seventh and final chapter of the online fsck documentation, where we talk about future functionality that can tie in with the functionality provided by the online fsck patchset. Signed-off-by: Darrick J. Wong <djwong@kernel.org> Reviewed-by: Dave Chinner <dchinner@redhat.com>
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Darrick J. Wong
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@ -5103,3 +5103,213 @@ mapping ioctl to map the recorded media errors back to metadata structures
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and report what has been lost.
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For media errors in blocks owned by files, parent pointers can be used to
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construct file paths from inode numbers for user-friendly reporting.
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7. Conclusion and Future Work
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=============================
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It is hoped that the reader of this document has followed the designs laid out
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in this document and now has some familiarity with how XFS performs online
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rebuilding of its metadata indices, and how filesystem users can interact with
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that functionality.
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Although the scope of this work is daunting, it is hoped that this guide will
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make it easier for code readers to understand what has been built, for whom it
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has been built, and why.
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Please feel free to contact the XFS mailing list with questions.
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FIEXCHANGE_RANGE
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----------------
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As discussed earlier, a second frontend to the atomic extent swap mechanism is
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a new ioctl call that userspace programs can use to commit updates to files
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atomically.
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This frontend has been out for review for several years now, though the
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necessary refinements to online repair and lack of customer demand mean that
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the proposal has not been pushed very hard.
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Extent Swapping with Regular User Files
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```````````````````````````````````````
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As mentioned earlier, XFS has long had the ability to swap extents between
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files, which is used almost exclusively by ``xfs_fsr`` to defragment files.
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The earliest form of this was the fork swap mechanism, where the entire
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contents of data forks could be exchanged between two files by exchanging the
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raw bytes in each inode fork's immediate area.
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When XFS v5 came along with self-describing metadata, this old mechanism grew
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some log support to continue rewriting the owner fields of BMBT blocks during
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log recovery.
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When the reverse mapping btree was later added to XFS, the only way to maintain
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the consistency of the fork mappings with the reverse mapping index was to
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develop an iterative mechanism that used deferred bmap and rmap operations to
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swap mappings one at a time.
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This mechanism is identical to steps 2-3 from the procedure above except for
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the new tracking items, because the atomic extent swap mechanism is an
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iteration of an existing mechanism and not something totally novel.
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For the narrow case of file defragmentation, the file contents must be
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identical, so the recovery guarantees are not much of a gain.
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Atomic extent swapping is much more flexible than the existing swapext
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implementations because it can guarantee that the caller never sees a mix of
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old and new contents even after a crash, and it can operate on two arbitrary
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file fork ranges.
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The extra flexibility enables several new use cases:
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- **Atomic commit of file writes**: A userspace process opens a file that it
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wants to update.
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Next, it opens a temporary file and calls the file clone operation to reflink
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the first file's contents into the temporary file.
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Writes to the original file should instead be written to the temporary file.
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Finally, the process calls the atomic extent swap system call
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(``FIEXCHANGE_RANGE``) to exchange the file contents, thereby committing all
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of the updates to the original file, or none of them.
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.. _swapext_if_unchanged:
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- **Transactional file updates**: The same mechanism as above, but the caller
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only wants the commit to occur if the original file's contents have not
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changed.
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To make this happen, the calling process snapshots the file modification and
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change timestamps of the original file before reflinking its data to the
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temporary file.
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When the program is ready to commit the changes, it passes the timestamps
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into the kernel as arguments to the atomic extent swap system call.
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The kernel only commits the changes if the provided timestamps match the
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original file.
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- **Emulation of atomic block device writes**: Export a block device with a
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logical sector size matching the filesystem block size to force all writes
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to be aligned to the filesystem block size.
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Stage all writes to a temporary file, and when that is complete, call the
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atomic extent swap system call with a flag to indicate that holes in the
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temporary file should be ignored.
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This emulates an atomic device write in software, and can support arbitrary
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scattered writes.
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Vectorized Scrub
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----------------
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As it turns out, the :ref:`refactoring <scrubrepair>` of repair items mentioned
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earlier was a catalyst for enabling a vectorized scrub system call.
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Since 2018, the cost of making a kernel call has increased considerably on some
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systems to mitigate the effects of speculative execution attacks.
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This incentivizes program authors to make as few system calls as possible to
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reduce the number of times an execution path crosses a security boundary.
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With vectorized scrub, userspace pushes to the kernel the identity of a
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filesystem object, a list of scrub types to run against that object, and a
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simple representation of the data dependencies between the selected scrub
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types.
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The kernel executes as much of the caller's plan as it can until it hits a
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dependency that cannot be satisfied due to a corruption, and tells userspace
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how much was accomplished.
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It is hoped that ``io_uring`` will pick up enough of this functionality that
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online fsck can use that instead of adding a separate vectored scrub system
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call to XFS.
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The relevant patchsets are the
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`kernel vectorized scrub
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<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=vectorized-scrub>`_
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and
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`userspace vectorized scrub
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<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=vectorized-scrub>`_
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series.
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Quality of Service Targets for Scrub
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------------------------------------
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One serious shortcoming of the online fsck code is that the amount of time that
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it can spend in the kernel holding resource locks is basically unbounded.
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Userspace is allowed to send a fatal signal to the process which will cause
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``xfs_scrub`` to exit when it reaches a good stopping point, but there's no way
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for userspace to provide a time budget to the kernel.
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Given that the scrub codebase has helpers to detect fatal signals, it shouldn't
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be too much work to allow userspace to specify a timeout for a scrub/repair
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operation and abort the operation if it exceeds budget.
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However, most repair functions have the property that once they begin to touch
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ondisk metadata, the operation cannot be cancelled cleanly, after which a QoS
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timeout is no longer useful.
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Defragmenting Free Space
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------------------------
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Over the years, many XFS users have requested the creation of a program to
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clear a portion of the physical storage underlying a filesystem so that it
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becomes a contiguous chunk of free space.
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Call this free space defragmenter ``clearspace`` for short.
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The first piece the ``clearspace`` program needs is the ability to read the
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reverse mapping index from userspace.
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This already exists in the form of the ``FS_IOC_GETFSMAP`` ioctl.
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The second piece it needs is a new fallocate mode
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(``FALLOC_FL_MAP_FREE_SPACE``) that allocates the free space in a region and
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maps it to a file.
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Call this file the "space collector" file.
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The third piece is the ability to force an online repair.
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To clear all the metadata out of a portion of physical storage, clearspace
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uses the new fallocate map-freespace call to map any free space in that region
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to the space collector file.
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Next, clearspace finds all metadata blocks in that region by way of
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``GETFSMAP`` and issues forced repair requests on the data structure.
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This often results in the metadata being rebuilt somewhere that is not being
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cleared.
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After each relocation, clearspace calls the "map free space" function again to
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collect any newly freed space in the region being cleared.
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To clear all the file data out of a portion of the physical storage, clearspace
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uses the FSMAP information to find relevant file data blocks.
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Having identified a good target, it uses the ``FICLONERANGE`` call on that part
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of the file to try to share the physical space with a dummy file.
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Cloning the extent means that the original owners cannot overwrite the
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contents; any changes will be written somewhere else via copy-on-write.
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Clearspace makes its own copy of the frozen extent in an area that is not being
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cleared, and uses ``FIEDEUPRANGE`` (or the :ref:`atomic extent swap
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<swapext_if_unchanged>` feature) to change the target file's data extent
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mapping away from the area being cleared.
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When all other mappings have been moved, clearspace reflinks the space into the
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space collector file so that it becomes unavailable.
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There are further optimizations that could apply to the above algorithm.
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To clear a piece of physical storage that has a high sharing factor, it is
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strongly desirable to retain this sharing factor.
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In fact, these extents should be moved first to maximize sharing factor after
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the operation completes.
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To make this work smoothly, clearspace needs a new ioctl
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(``FS_IOC_GETREFCOUNTS``) to report reference count information to userspace.
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With the refcount information exposed, clearspace can quickly find the longest,
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most shared data extents in the filesystem, and target them first.
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**Future Work Question**: How might the filesystem move inode chunks?
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*Answer*: To move inode chunks, Dave Chinner constructed a prototype program
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that creates a new file with the old contents and then locklessly runs around
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the filesystem updating directory entries.
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The operation cannot complete if the filesystem goes down.
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That problem isn't totally insurmountable: create an inode remapping table
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hidden behind a jump label, and a log item that tracks the kernel walking the
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filesystem to update directory entries.
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The trouble is, the kernel can't do anything about open files, since it cannot
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revoke them.
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**Future Work Question**: Can static keys be used to minimize the cost of
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supporting ``revoke()`` on XFS files?
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*Answer*: Yes.
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Until the first revocation, the bailout code need not be in the call path at
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all.
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The relevant patchsets are the
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`kernel freespace defrag
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<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfs-linux.git/log/?h=defrag-freespace>`_
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and
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`userspace freespace defrag
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<https://git.kernel.org/pub/scm/linux/kernel/git/djwong/xfsprogs-dev.git/log/?h=defrag-freespace>`_
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series.
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Shrinking Filesystems
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---------------------
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Removing the end of the filesystem ought to be a simple matter of evacuating
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the data and metadata at the end of the filesystem, and handing the freed space
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to the shrink code.
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That requires an evacuation of the space at end of the filesystem, which is a
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use of free space defragmentation!
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