6.S081 2020 Lecture 16: Linux ext3 crash recovery

This lecture: a case study of logging in Linux's ext3 file system
  with an eye to performance and real-world design details

Why do we care about logging?
  it's a successful and important idea
  it magically adds crash recovery to almost any storage system
  widely used
    many forms of persistent storage (DBs, file systems, &c)
    fault tolerance in distributed systems
  lots to chew on for performance, correctness, flexibility

Review of xv6 logging
  [ on-disk FS, in-memory cache, on-disk log ]
  each system call is a transaction: begin_op(), writes, end_op()
    initially writes affect only cached blocks in memory
  the overall goal:
    make a system call's writes atomic w.r.t. crashes
    apply all of the writes to the FS, or none of them
    e.g. mark block as free AND erase block # from i-node
  at end of system call:
    1. write all modified blocks to log on disk
    2. write block #s to log header on disk -- the commit point
    3. install modified blocks in "home" locations in FS on disk
    4. zero out log header on disk
       so logged blocks from next xaction don't appear committed
  if crash (e.g. computer loses power, but disk OK):
    reboot, then run log recovery code (before running any processes!)
    if log header contains block numbers
      replay all the logged writes to home locations
      zero out the log header
  write-ahead rule:
    don't start home FS writes until all writes committed in log on disk
    otherwise, if crash, recovery cannot complete the xaction
  freeing rule:
    don't erase/overwrite log until all home FS writes are done

What's wrong with xv6's logging?
  it is slow!
  each syscall has to wait for all disk I/O-- "synchronous update"
  new syscalls have to wait for commit to finish
    prevents concurrent execution of multiple processes
  every block written twice to disk, once to log, once to FS

Linux's ext3 design
  case study of what it takes to make logging efficient
  ext3 adds a log to ext2, a previous log-less file system
  (the reading calls it ext2, but now known as ext3)
  (the reading's "journal" means the same thing as "log")

ext3 data structures:
  in memory:
    write-back block cache
    per-transaction info:
      sequence #
      set of block #s to be logged
      set of outstanding "handles" -- one per syscall
  on disk:
    FS
    circular log

what's in the ext3 log (== journal)?
  log superblock: log offset and seq # of earliest valid transaction
    (this is not the FS superblock; it's a block at start of log file)
  descriptor blocks: magic, seq, block #s
  data blocks (as described by descriptor)
  commit blocks: magic, seq
 
the log can hold multiple transactions
  |super: offset+seq #| ... |Descriptor 4|...blocks...|Commit 4| |Descriptor 5|

how does ext3 get good performance?
  asynchronous disk update
  batching
  concurrency

asynchronous disk updates
  system calls only modify blocks in cache
    they do not update the disk (and thus don't wait for disk writes)
  why does async update help performance?
    1. system calls return quickly
    2. I/O concurrency; app can do other things while disk writes
    3. allows batching

asynchronous disk updates affect ext3's crash semantics
  a crash may "forget" the last few completed system calls
  this is a significant different from xv6 and many non-logging file systems
  an application may tell a user "I finished writing the file"
    but a crash may cause the file to disappear after crash+recovery
  careful applications need to be aware (e.g. DB or mail server or editor)
  fsync(fd) forces updates to log (and commit) before returning
    databases, text editors, &c call fsync()

ext3 batching
  ext3 commits every few seconds
    so each transaction typically includes many syscalls
  why does batching help performance?
    1. spread fixed transaction costs over many system calls
       fixed costs: descriptor and commit blocks, seek/rotate, syscall barrier
    2. "write absorbtion"
       many syscalls in the batch may modify the same block (bitmap, dirent)
       thus one disk write for many syscalls' updates
    3. disk scheduling: sort writes for fewer seeks, sequential runs

ext3 allows two kinds of concurrency
  1. multiple system calls can add to the current transaction
  2. multiple transactions may be at various stages:
     * one "open" transaction that's accepting new syscalls 
     * committing transactions doing their log writes
     * committed transactions writing to home locations in FS on disk
     * old transactions being freed
  why does concurrency help performance?
    CPU concurrency: many processes can use FS in parallel on multi-core
    I/O concurrency: processes can operate on FS in cache while
      previous transactions are writing the disk

ext3 sys call code:
  sys_unlink() {
    h = start()
    get(h, block #) ...
    modify the block(s) in the cache
    stop(h)
  }
  start():
    tells the logging system about a new system call
      can't commit until all start()ed system calls have called stop()
    start() can block this sys call if needed
  get():
    tells logging system we'll modify cached block
      added to list of blocks to be logged in current transaction
  stop():
    stop() does *not* cause a commit
    transaction can commit iff all included syscalls have called stop()

committing a transaction to disk
  1. block new syscalls
  2. wait for in-progress syscalls to stop()
  3. open a new transaction, unblock new syscalls
  4. write descriptor to on-disk log w/ list of block #s
  5. write modified blocks from cache to on-disk log
  6. wait for descriptor and data log disk writes to finish
  7. write the commit record to on-disk log
  8. wait for the commit disk write to finish
     this is the "commit point"
  9. now modified blocks allowed to go to homes on disk (but not forced)

when can ext3 re-use log space?
  log is circular: SB T4 T1 T2 T3
  T2's log space can be reused if:
    T2 has committed
    and T2's blocks have all been written to home locations in FS on disk
      ==> no longer any need to replay from log after a crash
    and all transactions prior to T2 have been freed in the log
  freeing writes on-disk log superblock with
    offset/seq of resulting oldest transaction in log

what if a crash?
  crash causes RAM (and thus cached disk blocks) to be lost
    but disk is assumed to be OK
  crash may interrupt writing last xaction to on-disk log
  so on-disk log may have a bunch of complete xactions, then one partial
  may also have written some of block cache to disk
    but only for fully committed xactions, not partial last one

how ext3 recovery works
  SB T8 xT5 T6 T7
    SB points to T6
    T5 is free, T8 has overwritten start of T5
  0. reboot with intact disk
  1. look in log "superblock" for offset and seq# of oldest transaction
  2. find the end of the log
     scan until bad magic (missing commit) or unexpected seq # (old entry)
     go back to last valid commit block, to ignore final partial transaction
       in case the crash occured midway through writing a transaction
  3. replay all blocks through last valid commit, in log order

what if block after last valid commit block looks like a log descriptor?
  i.e. looks like the start of a new transaction?
  perhaps a descriptor block left over from previous use of log?
    no: seq # will be too low
  perhaps some file data happens to look like a descriptor?
    ext3 prevents this from happening
    ext3 replaces magic number in data blocks in journal with zero
    and sets a flag for that block in the descriptor

what if another crash during log replay?
  after reboot, recovery will replay exactly the same log writes

that was the straightforward part of ext3.
  there are also a bunch of tricky details!

why does ext3 delay start of T2's syscalls until all of T1's syscalls finish?
  i.e. why this:
    T1: |-syscalls-|
    T2:             |-syscalls-|
              ---time--->
  this barrier sacrifices some performance
  example problem that this barrier prevents:
    T1: |-create(x)-|
    T2:        |-unlink(y)-|
                       X crash
              ---time--->
  if we allow unlink to start before create finishes
    unlink(y) free y's i-node -- i.e. marks it free in block cache
    create(x) may allocate that same i-node -- i.e. read unlink's write
  if T1 commits, but crash before T2 commits, then
    y will exist (the unlink did not commit)
    x will exist but use the same i-node as y!
  we can't let T1 write T2's updates to the log!
    so: don't start T2's syscalls until T1's syscalls are done
  (note ext3's copy-on-write doesn't fix this: it gives T1
   a copy of all blocks as of the end of T1, but the unlink
   executed before T1 ended.)
  (note we can't give T2 its own copies of the blocks, separate from
   T1's copies, since generally we do want later system calls to
   see the effects of earlier system calls.)
    
how about this:
  T1: |--create(d/f)--|
  T2:                   |--create(d/g)--|
  both create()s write d's directory content block
  suppose T1 commits (writes log on disk) *after* T2's create()
  we can't let T1 write T2's update to the log!
    then a crash plus replay of would replay some but not all of T2
    e.g. make the directory entry d/g but without initializing g's i-node
  ext3's solution:
    give T1 a private copy of the block cache as it existed when T1 closed
    T1 commits from this snapshot of the cache
    it's efficient using copy-on-write
    the copies allow syscalls in T2 to proceed while T1 is committing

so far we've been talking about EXT3's "journaled data" mode
  in which file content blocks are written to log
  as well as meta-data (i-nodes, directory content, free bitmaps)
  so file content is included in atomicity guarantee
  e.g. a write() that grows a file results in these log entries:
    [descriptor], i-node, block bitmap, content, [commit]

logging file content is slow, every data block written twice
  data is usually much bigger than meta-data
  can we entirely omit file content from the log?
  if we did, when would we write file content to the FS?

can we write file content blocks at any time at all?
  no: if metadata is committed first, crash may leave i-node pointing
      to blocks from someone else's deleted file, with previous
      file's content!

ext3 "ordered data" mode:
  write file content blocks to disk *before* commiting log
  (and don't write file content to the log)
  thus, for a write():
    write file content blocks to FS on disk, wait, then
    commit updated i-node (w/ block number and increased size) to the log
  if crash after data write, but before commit:
    block on disk has new data
    but not visible, since i-node size and block list not updated
    no metadata inconsistencies
      neither i-node nor free bitmap were updated, so blocks still free
  most people use ext3 ordered data mode
    it's much faster: avoids writing data twice

why does it make sense for the log to ignore file content?
  the real point of the log is to protect the FS internal invariants
    e.g. a block cannot be both "free" and used in an i-node
  in order to have fast recovery after crash + reboot.
  from this point of view, file content is not relevant.
  in any case, the file system doesn't know enough about application
    writes to understand which groups of writes need to be atomic.
  instead, applications can call fsync() to ensure data is
    safe on the disk.

correctness challenges w/ ordered data mode:
  A. rmdir, re-use directory content block for write() to some file,
       crash before rmdir or write committed
     after recovery, as if rmdir never happened,
       but directory content block has been overwritten!
     fix: no re-use of freed block until freeing syscall committed
  B. rmdir, commit, re-use block in file, ordered file write, commit,
       crash+recover, replay rmdir
     but no replay of file content write!
     file is left w/ directory content e.g. . and ..
     fix: put "revoke" records into log, prevent log replay of a given block
  note: both problems due to changing the type of a block (content vs meta-data)
    so another solution might be to never change a block's type
	 
Summary of ext3 rules
  The basic write-ahead logging rule:
    Don't write meta-data block to on-disk FS until committed in on-disk log
  Wait for all syscalls in T1 to finish before starting T2
  Don't overwrite a block in buffer cache before it is in the log
  Don't free on-disk log transaction until all blocks have been written to FS
  Ordered data mode:
    Write datablock to FS before commit
    Don't reuse free block until freeing syscall is committed
    Don't replay revoked syscalls

another corner case: open fd and unlink
  open a file, then unlink it
  file is open, so unlink removes dir entry but doesn't free inode or blocks
  unlink commits (just removal of dir entry)
  crash
  log doesn't contain writes that free the i-node or blocks
  inode and blocks not on free list, also not reachably by any name
    will never be freed! oops
  solution: add inode to linked list starting from FS superblock
    commit that along with remove of dir ent
  recovery looks at that list, completes deletions

what if there's not enough free log space for a transaction's blocks?
  free oldest transaction(s)
  i.e. install its writes in the on-disk FS

what if so many syscalls have started that the entire log isn't
  big enough to hold their writes, even if all older transactions
  have been freed?
  (unlikely, since each syscall generates few writes compared to log size)
  syscall passes max number of blocks needed to start()
  start() waits if total for current transaction is too high
    and works on freeing old transactions

another (complex) reason for reservations:
  T1 writes block 17
  T1 commits
  T1 finishes writing its log and commit record
     but has not yet written block 17 to home location
  T2 writes block 17 in the cache
     ext3 does *not* make a copy of block 17
     ext3 only copies if needed to write T1's log
  a T2 syscall is executing, does a write for which no log space
  can ext3 install T1's blocks to home locations on disk,
    and free T1's log space
  no: the cache no longer holds T1's version of block 17
    (ext3 could read the block from the log, but it doesn't)
  reservations detect the potential log space exhaustion,
    prevent that T2 syscall from starting

checksums
  recall: transaction's log blocks must be on disk before writing commit block
    ext3 waits for disk to say "done" before starting commit block write
  risk: disks usually have write caches and re-order writes, for performance
    sometimes hard to turn off (the disk lies)
    people often leave re-ordering enabled for speed, out of ignorance
    bad news if disk writes commit block before the rest of the transaction
  solution: commit block contains checksum of all data blocks
    on recovery: compute checksum of datablocks
	if matches checksum in commit block: install transaction
    if no match: don't install transaction
  ext4 has log checksumming, but ext3 does not

take-home points about logging and ext3
  logs make multi-step disk updates atomic w.r.t. crashes
  write-ahead rule: first commit in log, then write to disk
  recovery relies on this rule -- and is fast
  batching and concurrency give ext3 much better performance than xv6
    at the cost of many corner cases to be dealt with

References:
Stephen Tweedie 2000 talk transcript "EXT3, Journaling Filesystem"
  http://olstrans.sourceforge.net/release/OLS2000-ext3/OLS2000-ext3.html
log format details
  https://ext4.wiki.kernel.org/index.php/Ext4_Disk_Layout
https://www.usenix.org/system/files/login/articles/04_tso_018-021_final.pdf