From: Pedro F. Giffuni <giffunip@...>
Subject: [Tux3] Comparison to Hammer fs design
Date: Jul 25, 12:22 pm 2008

Hi;

The announcement of yet another filesystem:

http://lkml.org/lkml/2008/7/23/257

led to some comments about hammer fs:

http://tux3.org/pipermail/tux3/2008-July/000006.html

enjoy,

     Pedro.

From: Matthew Dillon <dillon@...>
Subject: Re: [Tux3] Comparison to Hammer fs design
Date: Jul 25, 2:53 pm 2008

:Hi;
:
:The announcement of yet another filesystem:
:
:http://lkml.org/lkml/2008/7/23/257
:
:led to some comments about hammer fs:
:
:http://tux3.org/pipermail/tux3/2008-July/000006.html
:
:enjoy,
:
:     Pedro.

    Those are interesting comments.   I think I found Daniel's email address
    so I am adding him to the To:   Dan, feel free to post this on your Tux
    groups if you want.

    I did consider multiple-parentage...  that is the ability to have a
    writable snapshot that 'forks' the filesystem history.  It would be
    an ultra cool feature to have but I couldn't fit it into the B-Tree
    model I was using.  Explicit snapshotting would be needed to make it
    work, and the snapshot id would have to be made part of the B-Tree key,
    which is fine.  HAMMER is based around implicit snapshots (being able
    to do an as-of historical lookup without having explicitly snapshotted
    bits of the history).  I would caution against using B-Tree iterations
    in historical lookups, B-Trees are only fast if you can pretty much
    zero in on the exact element you want as part of the primary search
    mechanic.  Once you have to iterate or chain performance goes out the
    window.  I know this because there are still two places where HAMMER
    sometimes has to iterate due to the inexact nature of an as-of lookup.
    Multiple-parentage almost certainly require two inequality searches,
    or one inequality search and an iteration.  A single timeline only
    requires one inequality search.

    I couldn't get away from having a delete_tid (the 'death version
    numbers').  I really tried :-)  There are many cases where one is only
    deleting, rather then overwriting.   Both numbers are then needed to
    be able to properly present a historical view of the filesystem.
    For example, when one is deleting a directory entry a snapshot after
    that deletion must show the directory entry gone.  Both numbers are
    also needed to be able to properly prune out unwanted historical data
    from the filesystem.  HAMMER's pruning algorithm (cleaning out old 
    historical data which is no longer desired) creates holes in the
    sequence so once you start pruning out unwanted historical data
    the delete_tid of a prior record will not match the create_tid of the
    following one (historically speaking).

    At one point I did collapse the holes, rematching the delete_tid with
    the create_tid of the following historical record, but I had to remove
    that code because it seriously complicated the mirroring implementation.
    I wanted each mirror to be able to have its own historical retention
    policy independant of the master.  e.g. so you could mirror to a backup
    system which would retain a longer and more coarse-grained history then
    the production system.

    --

    I also considered increasing the B-Tree fan-out to 256 but decided
    against it because insertions and deletions really bloated up the
    UNDO FIFO.   Frankly I'm still undecided as to whether that was a good
    idea, I would have prefered 256.  I can actually compile in 256 by
    changing a #define, but I released with 64 because I hit up against
    a number of performance issues:  bcopy() overheads for insertions
    and deletions in certain tests became very apparent.  Locking
    conflicts became a much more serious issue because I am using whole-node
    locks rather then element locks.  And, finally, the UNDO records got
    really bloated.  I would need to adjust the node locking and UNDO
    generation code significantly to remove the bottlenecks before I could
    go to a 256-element B-Tree node.

    HAMMER's B-Tree elements are probably huge compared to Tux3, and that's
    another limiting factor for the fan-out I can use.  My elements
    are 64 bytes each.  64x64 = 4K per B-Tree node.  I decided to go
    with fully expanded keys: 64 bit object id, 64 bit file-offset/db-key,
    64 bit create_tid, 64 bit delete_tid), plus a 64-bit storage offset and
    other auxillary info.  That's instead of using a radix-compressed key.
    Radix compressed keys would have doubled the complexity of the B-Tree
    code, particularly with the middle-of-tree pointer caching that
    HAMMER does.

    The historical access mechanic alone added major complexity to the
    B-Tree algorithms.  I will note here that B-Tree searches have a very
    high cpu overhead no matter how you twist it, and being able to cache
    cursors within the B-Tree is absolutely required if you want the
    filesystem to perform well.  If you always start searches at the root
    your cpu overhead will be horrendous... so plan on caching cursors
    from the get-go.

    If I were to do radix compression I would also want to go with a
    fully dynamic element size and fully dynamic fan-out in order to
    best-compress each B-Tree node.  Definitely food for thought.  I'd
    love to do something like that.  I think radix compression would
    remove much of the topological bloat the B-Tree creates verses using
    blockmaps, generally speaking.

    Space management is currently HAMMER's greatest weakness, but it only
    applies to small storage systems.  Several things in HAMMER's design
    are simply not condusive to a small storage.  The storage model is not
    fine-grained and (unless you are deleting a lot of stuff) needs
    reblocking to actually recover the freed space.  The flushing algorithms
    need around 100MB of UNDO FIFO space on-media to handle worst-case
    dependancies (mainly directory/file visibility issues on crash recovery),
    and the front-end caching of modifying operations, since they don't
    know exactly how much actual storage will be needed, need ~16MB of
    wiggle room to be able to estimate the on-media storage required to
    back the operations cached by the front-end.  Plus, on top of that,
    any sort of reblocking also needs some wiggle room to be able to clean
    out partially empty big-blocks and fill in new ones.

    On the flip side, the reblocker doesn't just de-fragment the filesystem,
    it is also intended to be used for expanding and contracting partitions,
    and adding removing volumes in the next release.  Definitely a
    multi-purpose utility.

    So I'm not actually being cavalier about it, its just that I had to
    make some major decisions on the algorithm design and I decided to
    weight the design more towards performance and large-storage, and
    small-storage suffered a bit.

    --

    In anycase, it sounds like Tux3 is using many similar ideas.  I think
    you are on the right track.  I will add one big note of caution, drawing
    from my experience implementing HAMMER, because I think you are going
    to hit a lot of the same issues.

    I spent 9 months designing HAMMER and 9 months implementing it.  During
    the course of implementing it I wound up throwing away probably 80% of
    the original design outright.  Amoung the major components I had to
    rewrite were the LOG mechanic (which ultimately became the meta-data
    UNDO FIFO), and the fine-grained storage mechanic (which ultimately
    became coarse-grained).  The UNDO FIFO was actually the saving grace,
    once that was implemented most of the writer-ordering dependancies went
    away (devolved into just flushing meta-data buffers after syncing the
    UNDO FIFO)... I suddenly had much, much more freedom in designing
    the other on-disk structures and algorithms.

    What I found while implementing HAMMER was that the on-disk topological
    design essentially dictated the extent of HAMMER's feature set AND
    most of its deficiencies (such as having to reblock to recover space),
    and the algorithms I chose often dictated other restrictions.  But the
    major restrictions came from the on-disk structures.

    Because of the necessarily tight integration between subsystems I found
    myself having to do major redesigns during the implementation phase.
    Fixing one subsystem created a cascade effect that required tweaking other
    subsystems.  Even adding new features, such as the mirroring, required
    significant changes in the B-Tree deadlock recovery code.  I couldn't get
    away from it.  Ultimately I chose to simplify some of the algorithms
    rather then have to go through another 4 months of rewriting.  All
    major designs are an exercise in making trade-offs in topology, feature
    set, algorithmic complexity, debuggability, robustness, etc.  The list
    goes on forever.

    Laters!

					-Matt
					Matthew Dillon 
					<dillon@backplane.com>


From: Daniel Phillips <phillips@...>
Subject: Re: [Tux3] Comparison to Hammer fs design
Date: Jul 25, 6:54 pm 2008

(resent after subscribing to the the kernel@crater)

On Friday 25 July 2008 11:53, Matthew Dillon wrote:
> 
> :Hi;
> :
> :The announcement of yet another filesystem:
> :
> :http://lkml.org/lkml/2008/7/23/257
> :
> :led to some comments about hammer fs:
> :
> :http://tux3.org/pipermail/tux3/2008-July/000006.html
> :
> :enjoy,
> :
> :     Pedro.
> 
>     Those are interesting comments.   I think I found Daniel's email address
>     so I am adding him to the To:   Dan, feel free to post this on your Tux
>     groups if you want.

How about a cross post?
 
>     I did consider multiple-parentage...  that is the ability to have a
>     writable snapshot that 'forks' the filesystem history.  It would be
>     an ultra cool feature to have but I couldn't fit it into the B-Tree
>     model I was using.  Explicit snapshotting would be needed to make it
>     work, and the snapshot id would have to be made part of the B-Tree key,
>     which is fine.  HAMMER is based around implicit snapshots (being able
>     to do an as-of historical lookup without having explicitly snapshotted
>     bits of the history).

Yes, that is the main difference indeed, essentially "log everything" vs
"commit" style versioning.  The main similarity is the lifespan oriented
version control at the btree leaves.

>     I would caution against using B-Tree iterations 
>     in historical lookups, B-Trees are only fast if you can pretty much
>     zero in on the exact element you want as part of the primary search
>     mechanic.  Once you have to iterate or chain performance goes out the
>     window.  I know this because there are still two places where HAMMER
>     sometimes has to iterate due to the inexact nature of an as-of lookup.
>     Multiple-parentage almost certainly require two inequality searches,
>     or one inequality search and an iteration.  A single timeline only
>     requires one inequality search.

Once I get down to the leaf level I binary search on logical address in
the case of a file index btree or on version in the case of an inode
table block, so this cost is still Log(N) with a small k.  For a
heavily versioned inode or file region this might sometimes result in
an overflow block or two that has to be linearly searched which is not
a big problem so long as it is a rare case, which it really ought to be
for the kinds of filesystem loads I have seen.  A common example of a
worst case is /var/log/messages, where the mtime and size are going to
change in pretty much every version, so if you have hourly snapshots
and hold them for three months it adds up to about 2200 16 byte inode
table attributes to record, about 8 inode table leaf blocks.  I really
do not see that as a problem.  If it goes up to 100 leaf blocks to
search then that could be a problem.

Usually, I will only be interested in the first and last of those
blocks, the first holding the rarely changing attributes including the
root of the file index btree and the last holding the most recent
incarnation of the mtime/size attribute.

The penultimate inode table index block tells me how many blocks a
given inode lives in because several blocks will have the same inum
key.  So the lookup algorithm for a massively versioned file becomes:
1) read the first inode table block holding that inum; 2) read the last
block with the same inum.  The latter operation only needs to consult
the immediate parent index block, which is locked in the page cache at
that point.

It will take a little care and attention to the inode format,
especially the inode header, to make sure that I can reliably do that
first/last probe optimization for the "head" version, but it does seem
worth the effort.  For starters I will just do a mindless linear walk
of an "overflow" inode and get fancy if that turns out to be a problem.

>     I couldn't get away from having a delete_tid (the 'death version
>     numbers').  I really tried :-)  There are many cases where one is only
>     deleting, rather then overwriting.

By far the most common case I would think.  But check out the versioned
pointer algorithms.  Surprisingly that just works, which is not exactly
obvious:

   Versioned pointers: a new method of representing snapshots
   http://lwn.net/Articles/288896/

I was originally planning to keep all versions of a truncate/rewrite
file in the same file index, but recently I realized that that is dumb
because there will never be any file data shared by a successor version
in that case.  So the thing to do is just create an entirely new
versioned file data attribute for each rewrite, bulking up the inode
table entry a little but greatly constraining the search for versions
to delete and reducing cache pressure by not loading unrelated version
data when traversing a file.

>     Both numbers are then needed to 
>     be able to properly present a historical view of the filesystem.
>     For example, when one is deleting a directory entry a snapshot after
>     that deletion must show the directory entry gone.

...which happens in Tux3 courtesy of the fact that the entire block
containing the dirent will have been versioned, with the new version
showing the entry gone.  Here is one of two places where I violate my
vow to avoid copying an entire block when only one data item in it
changes (the other being the atime table).  I rely on two things to
make this nice: 1) Most dirent changes will be logically logged and
only rolled up into the versioned file blocks when there are enough to
be reasonably sure that each changed directory block will be hit
numerous times in each rollup episode.  (Storing the directory blocks
in dirent-create order as in PHTree makes this very likely for mass
deletes.)  2) When we care about this is usually during a mass delete,
where most or all dirents in each directory file block are removed
before moving on to the next block.

>     Both numbers are 
>     also needed to be able to properly prune out unwanted historical data
>     from the filesystem.  HAMMER's pruning algorithm (cleaning out old  
>     historical data which is no longer desired) creates holes in the
>     sequence so once you start pruning out unwanted historical data
>     the delete_tid of a prior record will not match the create_tid of the
>     following one (historically speaking).

Again, check out the versioned pointer algorithms.  You can tell what
can be pruned just by consulting the version tree and the create_tids
(birth versions) for a particular logical address.  Maybe the hang is
that you do not organize the btrees by logical address (or inum in the
case of the inode table tree).  I thought you did but have not read
closely enough to be sure.

>     At one point I did collapse the holes, rematching the delete_tid with
>     the create_tid of the following historical record, but I had to remove
>     that code because it seriously complicated the mirroring implementation.
>     I wanted each mirror to be able to have its own historical retention
>     policy independant of the master.  e.g. so you could mirror to a backup
>     system which would retain a longer and more coarse-grained history then
>     the production system.

Fair enough.  I have an entirely different approach to what you call
mirroring and what I call delta replication.  (I reserve the term
mirroring to mean mindless duplication of physical media writes.)  This
method proved out well in ddsnap:

   http://phunq.net/ddtree?p=zumastor/.git;a=tree;h=fc5cb496fff10a2b03034fcf95122f5828149...

(Sorry about the massive URL, you can blame Linus for that;-)

What I do in ddsnap is compute all the blocks that differ between two
versions and apply those to a remote volume already holding the first
of the two versions, yielding a replica of the second version that is
logically but not physically identical.  The same idea works for a
versioned filesystem: compute all the leaf data that differs between
two versions, per inum, and apply the resulting delta to the
corresponding inums in the remote replica.  The main difference vs a
ddsnap volume delta is that not all of the changes are physical blocks,
there are also changed inode attributes, so the delta stream format
has to be elaborated accordingly.

>     I also considered increasing the B-Tree fan-out to 256 but decided
>     against it because insertions and deletions really bloated up the
>     UNDO FIFO.  Frankly I'm still undecided as to whether that was a good 
>     idea, I would have prefered 256.  I can actually compile in 256 by
>     changing a #define, but I released with 64 because I hit up against
>     a number of performance issues:  bcopy() overheads for insertions
>     and deletions in certain tests became very apparent.  Locking
>     conflicts became a much more serious issue because I am using whole-node
>     locks rather then element locks.  And, finally, the UNDO records got
>     really bloated.  I would need to adjust the node locking and UNDO
>     generation code significantly to remove the bottlenecks before I could
>     go to a 256-element B-Tree node.

I intend to log insertions and deletions logically, which keeps each
down to a few bytes until a btree rollup episode comes along to perform
updating of the btree nodes in bulk.  I am pretty sure this will work
for you as well, and you might want to check out that forward logging
trick.

That reminds me, I was concerned about the idea of UNDO records vs
REDO.  I hope I have this right: you delay acknowledging any write
transaction to the submitter until log commit has progressed beyond the
associated UNDO records.  Otherwise, if you acknowledge, crash, and
prune all UNDO changes, then you would end up with applications
believing that they had got things onto stable storage and be wrong
about that.  I have no doubt you did the right thing there, but it is
not obvious from your design documentation.

>     HAMMER's B-Tree elements are probably huge compared to Tux3, and that's
>     another limiting factor for the fan-out I can use.  My elements
>     are 64 bytes each.

Yes, I mostly have 16 byte elements and am working on getting most of
them down to 12 or 8.

>     64x64 = 4K per B-Tree node.  I decided to go 
>     with fully expanded keys: 64 bit object id, 64 bit file-offset/db-key,
>     64 bit create_tid, 64 bit delete_tid), plus a 64-bit storage offset and
>     other auxillary info.  That's instead of using a radix-compressed key.
>     Radix compressed keys would have doubled the complexity of the B-Tree
>     code, particularly with the middle-of-tree pointer caching that
>     HAMMER does.

I use two-stage lookup, or four stage if you count searches within
btree blocks.  This makes the search depth smaller in the case of small
files, and in the case of really huge files it adds depth exactly as
appropriate.  The index blocks end up cached in the page cache (the
buffer cache is just a flavor of page cache in recent Linux) so there
is little repeated descending through the btree indices.  Instead, most
of the probing is through the scarily fast radix tree code, which has
been lovingly optimized over the years to avoid cache line misses and
SMP lock contention.

I also received a proposal for a "fat" btree index from a collaborator
(Maciej) that included the file offset but I really did not like the...
fattening.  A two level btree yields less metadata overall which in my
estimation is more important than saving some bree probes.  The way
things work these days, falling down from level 1 cache to level 2 or
from level 2 to level 3 costs much more than executing some extra CPU
instructions.  So optimization strategy inexorably shifts away from
minimizing instructions towards minimizing cache misses.

>     The historical access mechanic alone added major complexity to the
>     B-Tree algorithms.  I will note here that B-Tree searches have a very
>     high cpu overhead no matter how you twist it, and being able to cache
>     cursors within the B-Tree is absolutely required if you want the
>     filesystem to perform well.  If you always start searches at the root
>     your cpu overhead will be horrendous... so plan on caching cursors
>     from the get-go.

Fortunately, we get that for free on Linux, courtesy of the page cache
radix trees :-)

I might eventually add some explicit cursor caching, but various
artists over the years have noticed that it does not make as much
difference as you might think.

>     If I were to do radix compression I would also want to go with a
>     fully dynamic element size and fully dynamic fan-out in order to
>     best-compress each B-Tree node.  Definitely food for thought.

Indeed.  But Linux is braindamaged about large block size, so there is
very strong motivation to stay within physical page size for the
immediate future.  Perhaps if I get around to a certain hack that has
been perenially delayed, that situation will improve:

   "Variable sized page objects"
   http://lwn.net/Articles/37795/

>     I'd love to do something like that.  I think radix compression would
>     remove much of the topological bloat the B-Tree creates verses using
>     blockmaps, generally speaking.

Topological bloat?

>     Space management is currently HAMMER's greatest weakness, but it only
>     applies to small storage systems.  Several things in HAMMER's design
>     are simply not condusive to a small storage.  The storage model is not
>     fine-grained and (unless you are deleting a lot of stuff) needs
>     reblocking to actually recover the freed space.  The flushing algorithms
>     need around 100MB of UNDO FIFO space on-media to handle worst-case
>     dependancies (mainly directory/file visibility issues on crash recovery),
>     and the front-end caching of modifying operations, since they don't
>     know exactly how much actual storage will be needed, need ~16MB of
>     wiggle room to be able to estimate the on-media storage required to
>     back the operations cached by the front-end.  Plus, on top of that,
>     any sort of reblocking also needs some wiggle room to be able to clean
>     out partially empty big-blocks and fill in new ones.

Ah, that is a very nice benefit of Tux3-style forward logging I forgot
to mention in the original post: transaction size is limited only by
available free space on the volume, because log commit blocks can be
anywhere.  Last night I dreamed up a log commit block variant where the
associated transaction data blocks can be physically discontiguous, to
make this assertion come true, shortly after reading Stephen Tweedie's
horror stories about what you have to do if you must work around not
being able to put an entire, consistent transaction onto stable media
in one go:

   http://olstrans.sourceforge.net/release/OLS2000-ext3/OLS2000-ext3.html

Most of the nasties he mentions just vanish if:

  a) File data is always committed atomically
  b) All commit transactions are full transactions

He did remind me (via time travel from year 2000) of some details I
should write into the design explicitly, for example, logging orphan
inodes that are unlinked while open, so they can be deleted on replay
after a crash.  Another nice application of forward logging, which
avoids the seek-happy linked list through the inode table that Ext3
does.

>     On the flip side, the reblocker doesn't just de-fragment the filesystem,
>     it is also intended to be used for expanding and contracting partitions,
>     and adding removing volumes in the next release.  Definitely a
>     multi-purpose utility.

Good point.  I expect Tux3 will eventually have a reblocker (aka
defragger).  There are some really nice things you can do, like:

  1) Set a new version so logical changes cease for the parent
     version.

  2) We want to bud off a given directory tree into its own volume,
     so start by deleting the subtree in the current version.  If
     any link counts in the directory tree remain nonzero in the
     current version then there are hard links into the subtree, so
     fail now and drop the new version.

  3) Reblock a given region of the inode table tree and all the files
     in it into one physically contiguous region of blocks

  4) Add a free tree and other once-per volume structures to the new
     home region.

  5) The new region is now entirely self contained and even keeps its
     version history.  At the volume manager level, map it to a new,
     sparse volume that just has a superblock in the zeroth extent and
     the new, mini filesystem at some higher logical address.  Remap
     the region in the original volume to empty space and add the
     empty space to the free tree.

  6) Optionally reblock the newly budded filesystem to the base of the
     new volume so utilties that do not not play well with sparse
     volumes do not do silly things.

>     So I'm not actually being cavalier about it, its just that I had to
>     make some major decisions on the algorithm design and I decided to
>     weight the design more towards performance and large-storage, and
>     small-storage suffered a bit.

Cavalier was a poor choice of words, the post was full of typos as well
so I am not proud of it.  You are solving a somewhat different problem
and you have code out now, which is a huge achievement.  Still I think
you can iteratively improve your design using some of the techniques I
have stumbled upon.  There are probably some nice tricks I can get from
your code base too once I delve into it.

>     In anycase, it sounds like Tux3 is using many similar ideas.  I think
>     you are on the right track.

Thankyou very much.  I think you are on the right track too, which you
have a rather concrete way of proving.

>     I will add one big note of caution, drawing 
>     from my experience implementing HAMMER, because I think you are going
>     to hit a lot of the same issues.
> 
>     I spent 9 months designing HAMMER and 9 months implementing it.  During
>     the course of implementing it I wound up throwing away probably 80% of
>     the original design outright.  Amoung the major components I had to
>     rewrite were the LOG mechanic (which ultimately became the meta-data
>     UNDO FIFO), and the fine-grained storage mechanic (which ultimately
>     became coarse-grained).  The UNDO FIFO was actually the saving grace,
>     once that was implemented most of the writer-ordering dependancies went
>     away (devolved into just flushing meta-data buffers after syncing the
>     UNDO FIFO)... I suddenly had much, much more freedom in designing
>     the other on-disk structures and algorithms.
>
>     What I found while implementing HAMMER was that the on-disk topological
>     design essentially dictated the extent of HAMMER's feature set AND
>     most of its deficiencies (such as having to reblock to recover space),
>     and the algorithms I chose often dictated other restrictions.  But the
>     major restrictions came from the on-disk structures.
> 
>     Because of the necessarily tight integration between subsystems I found
>     myself having to do major redesigns during the implementation phase.
>     Fixing one subsystem created a cascade effect that required tweaking other
>     subsystems.  Even adding new features, such as the mirroring, required
>     significant changes in the B-Tree deadlock recovery code.  I couldn't get
>     away from it.  Ultimately I chose to simplify some of the algorithms
>     rather then have to go through another 4 months of rewriting.  All
>     major designs are an exercise in making trade-offs in topology, feature
>     set, algorithmic complexity, debuggability, robustness, etc.  The list
>     goes on forever.
> 

The big ahas! that eliminated much of the complexity in the Tux3 design
were:

  * Forward logging - just say no to incomplete transactions
  * Version weaving - just say no to recursive copy on write

Essentially I have been designing Tux3 for ten years now and working
seriously on the simplifying elements for the last three years or so,
either entirely on paper or in related work like ddsnap and LVM3.

One of the nice things about moving on from design to implementation of
Tux3 is that I can now background the LVM3 design process and see what
Tux3 really wants from it.  I am determined to match every checkbox
volume management feature of ZFS as efficiently or more efficiently,
without violating the traditional layering between filesystem and
block device, and without making LVM3 a Tux3-private invention.

>     Laters!

Hopefully not too much later.  I find this dialog very fruitful.  I just
wish such dialog would occur more often at the design/development stage
in Linux and other open source work instead of each group obsessively
ignoring all "competing" designs and putting huge energy into chatting
away about the numerous bugs that arise from rushing their design or
ignoring the teachings of history.

Regards,

Daniel


From: Matthew Dillon <dillon@...>
Subject: Re: [Tux3] Comparison to Hammer fs design
Date: Jul 25, 10:02 pm 2008

:>     so I am adding him to the To:   Dan, feel free to post this on your Tux
:>     groups if you want.
:
:How about a cross-post?

    I don't think it will work, only subscribers can post to the DFly groups,
    but we'll muddle through it :-)  I will include the whole of the previous
    posting so the DFly groups see the whole thing, if you continue to get
    bounces.

    I believe I have successfully added you as an 'alias address' to the
    DragonFly kernel list so you shouldn't get bounced if you Cc it now.
 
:Yes, that is the main difference indeed, essentially "log everything" vs
:"commit" style versioning.  The main similarity is the lifespan oriented
:version control at the btree leaves.

    Reading this and a little more that you describe later let me make
    sure I understand the forward-logging methodology you are using.
    You would have multiple individually-tracked transactions in
    progress due to parallelism in operations initiated by userland and each
    would be considered committed when the forward-log logs the completion
    of that particular operation?

    If the forward log entries are not (all) cached in-memory that would mean
    that accesses to the filesystem would have to be run against the log
    first (scanning backwards), and then through to the B-Tree?  You
    would solve the need for having an atomic commit ('flush groups' in
    HAMMER), but it sounds like the algorithmic complexity would be
    very high for accessing the log.

    And even though you wouldn't have to group transactions into larger
    commits the crash recovery code would still have to implement those
    algorithms to resolve directory and file visibility issues.  The problem
    with namespace visibility is that it is possible to create a virtually
    unending chain of separate but inter-dependant transactions which either
    all must go, or none of them.  e.g. creating a, a/b, a/b/c, a/b/x, a/b/c/d,
    etc etc.  At some point you have to be able to commit so the whole mess
    does not get undone by a crash, and many completed mini-transactions
    (file or directory creates) actually cannot be considered complete until
    their governing parent directories (when creating) or children (when
    deleting) have been committed.  The problem can become very complex.

    Last question here:  Your are forward-logging high level operations.
    You are also going to have to log meta-data (actual B-Tree manipulation)
    commits in order to recover from a crash while making B-Tree
    modifications.  Correct?  So your crash recovery code will have to handle
    both meta-data undo and completed and partially completed transactions.
    And there will have to be a tie-in between the meta-data commits and
    the transactions so you know which ones have to be replayed.  That
    sounds fairly hairy.  Have you figured out how you are doing to do that?

:Once I get down to the leaf level I binary search on logical address in
:the case of a file index btree or on version in the case of an inode
:table block, so this cost is still Log(N) with a small k.  For a
:heavily versioned inode or file region this might sometimes result in
:an overflow block or two that has to be linearly searched which is not
:a big problem so long as it is a rare case, which it really ought to be
:for the kinds of filesystem loads I have seen.  A common example of a
:worst case is /var/log/messages, where the mtime and size are going to
:change in pretty much every version, so if you have hourly snapshots
:and hold them for three months it adds up to about 2200 16 byte inode
:table attributes to record, about 8 inode table leaf blocks.  I really
:do not see that as a problem.  If it goes up to 100 leaf blocks to
:search then that could be a problem.
    
    I think you can get away with this as long as you don't have too many
    snapshots, and even if you do I noticed with HAMMER that only a small
    percentage of inodes have a large number of versions associated with
    them from normal production operation.  /var/log/messages
    is an excellent example of that.  Log files were effected the most though
    I also noticed that very large files also wind up with multiple versions
    of the inode, such as when writing out a terrabyte-sized file. 

    Even with a direct bypass for data blocks (but not their meta-data,
    clearly), HAMMER could only cache so much meta-data in memory before
    it had to finalize the topology and flush the inode out.  A
    terrabyte-sized file wound up with about 1000 copies of the inode
    prior to pruning (one had to be written out about every gigabyte or so).

:The penultimate inode table index block tells me how many blocks a
:given inode lives in because several blocks will have the same inum
:key.  So the lookup algorithm for a massively versioned file becomes:
:1) read the first inode table block holding that inum; 2) read the last
:block with the same inum.  The latter operation only needs to consult
:the immediate parent index block, which is locked in the page cache at
:that point.

    How are you dealing with expansion of the logical inode block(s) as
    new versions are added?  I'm assuming you are intending to pack the
    inodes on the media so e.g. a 128-byte inode would only take up
    128 bytes of media space in the best case.  Multiple inodes would be
    laid out next to each other logically (I assume), but since the physical
    blocks are larger they would also have to be laid out next to each
    other physically within any given backing block.  Now what happens
    when one has to be expanded?

    I'm sure this ties into the forward-log but even with the best algorithms
    you are going to hit limited cases where you have to expand the inode.
    Are you just copying the inode block(s) into a new physical allocation
    then?

:It will take a little care and attention to the inode format,
:especially the inode header, to make sure that I can reliably do that
:first/last probe optimization for the "head" version, but it does seem
:worth the effort.  For starters I will just do a mindless linear walk
:of an "overflow" inode and get fancy if that turns out to be a problem.
:

    Makes sense.  I was doing mindless linear walks of the B-Tree element
    arrays for most of HAMMER's implementation until it was stabilized well
    enough that I could switch to a binary search.  And I might change it
    again to start out with a heuristical index estimation.

:>     I couldn't get away from having a delete_tid (the 'death version
:>     numbers').  I really tried :-)  There are many cases where one is only
:>     deleting, rather then overwriting.
:
:By far the most common case I would think.  But check out the versioned
:pointer algorithms.  Surprisingly that just works, which is not exactly
:obvious:
:
:   Versioned pointers: a new method of representing snapshots
:   http://lwn.net/Articles/288896/

    Yes, it makes sense.  If the snapshot is explicitly taken then you
    can store direct references and chain, and you wouldn't need a delete
    id in that case.  From that article though the chain looks fairly
    linear.  Historical access could wind up being rather costly.

:I was originally planning to keep all versions of a truncate/rewrite
:file in the same file index, but recently I realized that that is dumb
:because there will never be any file data shared by a successor version
:in that case.  So the thing to do is just create an entirely new
:versioned file data attribute for each rewrite, bulking up the inode
:table entry a little but greatly constraining the search for versions
:to delete and reducing cache pressure by not loading unrelated version
:data when traversing a file.

    When truncating to 0 I would agree with your assessment.  For the 
    topology you are using you would definitely want to use different
    file data sets.  You also have to deal with truncations that are not
    to 0, that might be to the middle of a file.  Certainly not a
    common case, but still a case that has to be coded for.  If you treat
    truncation to 0 as a special case you will be adding considerable
    complexity to the algorithm.

    With HAMMER I chose to keep everything in one B-Tree, whether historical
    or current.  That way one set of algorithms handles both cases and code
    complexity is greatly reduced.  It isn't optimal... large amounts of
    built up history still slow things down (though in a bounded fashion).
    In that regard creating a separate topology for snapshots is a good
    idea.

:>     Both numbers are then needed to 
:>     be able to properly present a historical view of the filesystem.
:>     For example, when one is deleting a directory entry a snapshot after
:>     that deletion must show the directory entry gone.
:
:...which happens in Tux3 courtesy of the fact that the entire block
:containing the dirent will have been versioned, with the new version
:showing the entry gone.  Here is one of two places where I violate my
:vow to avoid copying an entire block when only one data item in it
:changes (the other being the atime table).  I rely on two things to
:make this nice: 1) Most dirent changes will be logically logged and
:only rolled up into the versioned file blocks when there are enough to
:be reasonably sure that each changed directory block will be hit
:numerous times in each rollup episode.  (Storing the directory blocks
:in dirent-create order as in PHTree makes this very likely for mass
:deletes.)  2) When we care about this is usually during a mass delete,
:where most or all dirents in each directory file block are removed
:before moving on to the next block.

    This could wind up being a sticky issue for your implementation.
    I like the concept of using the forward-log but if you actually have
    to do a version copy of the directory at all you will have to update the
    link (or some sort of) count for all the related inodes to keep track
    of inode visibility, and to determine when the inode can be freed and
    its storage space recovered.

    Directories in HAMMER are just B-Tree elements.  One element per
    directory-entry.  There are no directory blocks.   You may want to
    consider using a similar mechanism.  For one thing, it makes lookups
    utterly trivial... the file name is hashed and a lookup is performed
    based on the hash key, then B-Tree elements with the same hash key
    are iterated until a match is found (usually the first element is the
    match).  Also, when adding or removing directory entries only the
    directory inode's mtime field needs to be updated.  It's size does not.

    Your current directory block method could also represent a problem for
    your directory inodes... adding and removing directory entries causing
    size expansion or contraction could require rolling new versions
    of the directory inode to update the size field.  You can't hold too
    much in the forward-log without some seriously indexing.  Another
    case to consider along with terrabyte-sized files and log files.
    
:>     Both numbers are 
:>     also needed to be able to properly prune out unwanted historical data
:>     from the filesystem.  HAMMER's pruning algorithm (cleaning out old  
:>     historical data which is no longer desired) creates holes in the
:>     sequence so once you start pruning out unwanted historical data
:>     the delete_tid of a prior record will not match the create_tid of the
:>     following one (historically speaking).
:
:Again, check out the versioned pointer algorithms.  You can tell what
:can be pruned just by consulting the version tree and the create_tids
:(birth versions) for a particular logical address.  Maybe the hang is
:that you do not organize the btrees by logical address (or inum in the
:case of the inode table tree).  I thought you did but have not read
:closely enough to be sure.

    Yes, I see.  Can you explain the versioned pointer algorithm a bit more,
    it looks almost like a linear chain (say when someone is doing a daily
    snapshot).  It looks great for optimal access to the HEAD but it doesn't
    look very optimal if you want to dive into an old snapshot. 

    For informational purposes: HAMMER has one B-Tree, organized using
    a strict key comparison.  The key is made up of several fields which
    are compared in priority order:

	localization	- used to localize certain data types together and
			  to separate pseudo filesystems created within
			  the filesystem.
	obj_id		- the object id the record is associated with.
			  Basically the inode number the record is 
			  associated with.

	rec_type	- the type of record, e.g. INODE, DATA, SYMLINK-DATA,
			  DIRECTORY-ENTRY, etc.

	key		- e.g. file offset

	create_tid	- the creation transaction id

    Inodes are grouped together using the localization field so if you
    have a million inodes and are just stat()ing files, the stat
    information is localized relative to other inodes and doesn't have to
    skip file contents or data, resulting in highly localized accesses
    on the storage media.

    Beyond that the B-Tree is organized by inode number and file offset.
    In the case of a directory inode, the 'offset' is the hash key, so
    directory entries are organized by hash key (part of the hash key is
    an iterator to deal with namespace collisions).

    The structure seems to work well for both large and small files, for
    ls -lR (stat()ing everything in sight) style traversals as well as 
    tar-like traversals where the file contents for each file is read.

    The create_tid is the creation transaction id.  Historical accesses are
    always 'ASOF' a particular TID, and will access the highest create_tid
    that is still <= the ASOF TID.  The 'current' version of the filesystem
    uses an asof TID of 0xFFFFFFFFFFFFFFFF and hence accesses the highest
    create_tid.

    There is also a delete_tid which is used to filter out elements that
    were deleted prior to the ASOF TID.  There is currently an issue with
    HAMMER where the fact that the delete_tid is *not* part of the B-Tree
    compare can lead to iterations for strictly deleted elements, verses
    replaced elements which will have a new element with a higher create_tid
    that the B-Tree can seek to directly.  In fact, at one point I tried
    indexing on delete_tid instead of create_tid, but created one hell of a
    mess in that I had to physically *move* B-Tree elements being deleted
    instead of simply marking them as deleted.

:Fair enough.  I have an entirely different approach to what you call
:mirroring and what I call delta replication.  (I reserve the term
:mirroring to mean mindless duplication of physical media writes.)  This
:method proved out well in ddsnap:
:
:   http://phunq.net/ddtree?p=zumastor/.git;a=tree;h=fc5cb496fff10a2b03034fcf95122f5828149...
:
:(Sorry about the massive URL, you can blame Linus for that;-)
:
:What I do in ddsnap is compute all the blocks that differ between two
:versions and apply those to a remote volume already holding the first
:of the two versions, yielding a replica of the second version that is
:logically but not physically identical.  The same idea works for a
:versioned filesystem: compute all the leaf data that differs between
:two versions, per inum, and apply the resulting delta to the
:corresponding inums in the remote replica.  The main difference vs a
:ddsnap volume delta is that not all of the changes are physical blocks,
:there are also changed inode attributes, so the delta stream format
:has to be elaborated accordingly.

    Are you scanning the entire B-Tree to locate the differences?  It
    sounds you would have to as a fall-back, but that you could use the
    forward-log to quickly locate differences if the first version is
    fairly recent.

    HAMMER's mirroring basically works like this:  The master has synchronized
    up to transaction id C, the slave has synchronized up to transaction id A.
    The mirroring code does an optimized scan of the B-Tree to supply all
    B-Tree elements that have been modified between transaction A and C.
    Any number of mirroring targets can be synchronized to varying degrees,
    and can be out of date by varying amounts (even years, frankly).

    I decided to use the B-Tree to optimize the scan.  The B-Tree is
    scanned and any element with either a creation or deletion transaction
    id >= the slave's last synchronization point is then serialized and
    piped to the slave.

    To avoid having to scan the entire B-Tree I perform an optimization
    whereby the highest transaction id laid down at a leaf is propagated
    up the B-Tree all the way to the root.  This also occurs if a B-Tree
    node is physically deleted (due to pruning), even if no elements are
    actually present at the leaf within the transaction range.

    Thus the mirroring scan is able to skip any internal node (and its
    entire sub-tree) which has not been modified after the synchronization
    point, and is able to identify any leaf for which real, physical deletions
    have occured (in addition to logical deletions which simply set the
    delete_tid field in the B-Tree element) and pass along the key range
    and any remaining elements in that leaf for the target to do a
    comparative scan with.

    This allows incremental mirroring, multiple slaves, and also allows
    a mirror to go offline for months and then pop back online again
    and optimally pick up where it left off.  The incremental mirroring
    is important, the last thing I want to do is have to scan 2 billion
    B-Tree elements to do an incremental mirroring batch.

    The advantage is all the cool features I got by doing things that way.
    The disadvantage is that the highest transaction id must be propagated
    up the tree (though it isn't quite that bad because in HAMMER an entire
    flush group uses the same transaction id, so we aren't constantly
    repropagating new transaction id's up the same B-Tree nodes when
    flushing a particular flush group).

    You may want to consider something similar.  I think using the
    forward-log to optimize incremental mirroring operations is also fine
    as long as you are willing to take the 'hit' of having to scan (though
    not have to transfer) the entire B-Tree if the mirror is too far
    out of date.

:I intend to log insertions and deletions logically, which keeps each
:down to a few bytes until a btree rollup episode comes along to perform
:updating of the btree nodes in bulk.  I am pretty sure this will work
:for you as well, and you might want to check out that forward logging
:trick.

    Yes.  The reason I don't is because while it is really easy to lay down
    a forward-log, intergrating it into lookup operations (if you don't
    keep all the references cached in-memory) is really complex code.
    I mean, you can always scan it backwards linearly and that certainly
    is easy to do, but the filesystem's performance would be terrible.
    So you have to index the log somehow to allow lookups on it in
    reverse to occur reasonably optimally.  Have you figured out how you
    are going to do that?

    With HAMMER I have an in-memory cache and the on-disk B-Tree.  Just
    writing the merged lookup code (merging the in-memory cache with the
    on-disk B-Tree for the purposes of doing a lookup) was fairly complex.
    I would hate to have to do a three-way merge.... in-memory cache, on-disk
    log, AND on-disk B-Tree.  Yowzer!

:That reminds me, I was concerned about the idea of UNDO records vs
:REDO.  I hope I have this right: you delay acknowledging any write
:transaction to the submitter until log commit has progressed beyond the
:associated UNDO records.  Otherwise, if you acknowledge, crash, and
:prune all UNDO changes, then you would end up with applications
:believing that they had got things onto stable storage and be wrong
:about that.  I have no doubt you did the right thing there, but it is
:not obvious from your design documentation.

    The way it works is that HAMMER's frontend is almost entirely
    disconnected from the backend.  All meta-data operations are cached
    in-memory.  create, delete, append, truncate, rename, write...
    you name it.  Nothing the frontend does modifies any meta-data
    or mata-data buffers.

    The only sychronization point is fsync(), the filesystem syncer, and
    of course if too much in-memory cache is built up.  To improve
    performance, raw data blocks are not included... space for raw data
    writes is reserved by the frontend (without modifying the storage
    allocation layer) and those data buffers are written to disk by the
    frontend directly, just without any meta-data on-disk to reference
    them so who cares if you crash then!  It would be as if those blocks
    were never allocated in the first place.

    When the backend decides to flush the cached meta-data ops it breaks
    the meta-data ops into flush-groups whos dirty meta-data fits in the 
    system's buffer cache.  The meta-data ops are executed, building the
    UNDO, updating the B-Tree, allocating or finalizing the storage, and
    modifying the meta-data buffers in the buffer cache.  BUT the dirty
    meta-data buffers are locked into memory and NOT yet flushed to
    the media.  The UNDO *is* flushed to the media, so the flush groups
    can build a lot of UNDO up and flush it as they go if necessary.

    When the flush group has completed any remaining UNDO is flushed to the
    media, we wait for I/O to complete, the volume header's UNDO FIFO index
    is updated and written out, we wait for THAT I/O to complete, *then*
    the dirty meta-data buffers are flushed to the media.  The flushes at
    the end are done asynchronously (unless fsync()ing) and can overlap with
    the flushes done at the beginning of the next flush group.  So there
    are exactly two physical synchronization points for each flush.

    If a crash occurs at any point, upon remounting after a reboot
    HAMMER only needs to run the UNDOs to undo any partially committed
    meta-data.

    That's it.  It is fairly straight forward.

    For your forward-log approach you would have to log the operations 
    as they occur, which is fine since that can be cached in-memory. 
    However, you will still need to synchronize the log entries with
    a volume header update to update the log's index, so the recovery
    code knows how far the log extends.  Plus you would also have to log
    UNDO records when making actual changes to the permanent media
    structures (your B-Trees), which is independant of the forward-log
    entries you made representing high level filesystem operations, and
    would also have to lock the related meta-data in memory until the
    related log entries can be synchronized.  Then you would be able
    to flush the meta-data buffers.

    The forward-log approach is definitely more fine-grained, particularly
    for fsync() operations... those would go much faster using a forward
    log then the mechanism I use because only the forward-log would have
    to be synchronized (not the meta-data changes) to 'commit' the work.
    I like that, it would be a definite advantage for database operations.

:>     HAMMER's B-Tree elements are probably huge compared to Tux3, and that's
:>     another limiting factor for the fan-out I can use.  My elements
:>     are 64 bytes each.
:
:Yes, I mostly have 16 byte elements and am working on getting most of
:them down to 12 or 8.

    I don't know how you can make them that small.  I spent months just
    getting my elements down to 64 bytes.  The data reference alone for
    data blocks is 12 bytes (64 bit media storage offset and 32 bit length).

:>     64x64 = 4K per B-Tree node.  I decided to go 
:>     with fully expanded keys: 64 bit object id, 64 bit file-offset/db-key,
:>     64 bit create_tid, 64 bit delete_tid), plus a 64-bit storage offset and
:>     other auxillary info.  That's instead of using a radix-compressed key.
:>     Radix compressed keys would have doubled the complexity of the B-Tree
:>     code, particularly with the middle-of-tree pointer caching that
:>     HAMMER does.
:
:I use two-stage lookup, or four stage if you count searches within
:btree blocks.  This makes the search depth smaller in the case of small
:files, and in the case of really huge files it adds depth exactly as
:appropriate.  The index blocks end up cached in the page cache (the
:buffer cache is just a flavor of page cache in recent Linux) so there
:is little repeated descending through the btree indices.  Instead, most
:of the probing is through the scarily fast radix tree code, which has
:been lovingly optimized over the years to avoid cache line misses and
:SMP lock contention.
:
:I also received a proposal for a "fat" btree index from a collaborator
:(Maciej) that included the file offset but I really did not like the...
:fattening.  A two level btree yields less metadata overall which in my
:estimation is more important than saving some bree probes.  The way
:things work these days, falling down from level 1 cache to level 2 or
:from level 2 to level 3 costs much more than executing some extra CPU
:instructions.  So optimization strategy inexorably shifts away from
:minimizing instructions towards minimizing cache misses.

    The depth and complexity of your master B-Tree will definitely
    be smaller.  You are offloading both the file contents and snapshots.
    HAMMER incorporates both into a single global B-Tree.  This has been
    somewhat of a mixed bag for HAMMER.  On the plus side performance is
    very good for production operations (where the same filename is not
    being created over and over and over and over again and where the same
    file is not truncated over and over and over again).... locality of
    reference is extremely good in a single global B-Tree when accessing
    lots of small files or when accessing fewer larger files.  On the 
    negative side, performance drops noticeably if a lot of historical
    information (aka snapshots) has built up in the data set being
    accessed.  Even though the lookups themselves are extremely efficient,
    the access footprint (the size of the data set the system must cache)
    of the B-Tree becomes larger, sometimes much larger.

    I think the cpu overhead is going to be a bit worse then you are
    contemplating, but your access footprint (with regards to system memory
    use caching the meta-data) for non-historical accesses will be better.

    Are you going to use a B-Tree for the per-file layer or a blockmap?  And
    how are you going to optimize the storage for small files?  Are you
    just going to leave them in the log and not push a B-Tree for them?

:>     The historical access mechanic alone added major complexity to the
:>     B-Tree algorithms.  I will note here that B-Tree searches have a very
:>     high cpu overhead no matter how you twist it, and being able to cache
:>     cursors within the B-Tree is absolutely required if you want the
:>     filesystem to perform well.  If you always start searches at the root
:>     your cpu overhead will be horrendous... so plan on caching cursors
:>     from the get-go.
:
:Fortunately, we get that for free on Linux, courtesy of the page cache
:radix trees :-)
:
:I might eventually add some explicit cursor caching, but various
:artists over the years have noticed that it does not make as much
:difference as you might think.

     For uncacheable data sets the cpu overhead is almost irrelevant.

     But for cached data sets, watch out!  The cpu overhead of your B-Tree
     lookups is going to be 50% of your cpu time, with the other 50% being
     the memory copy or memory mapping operation and buffer cache operations.
     It is really horrendous.  When someone is read()ing or write()ing a
     large file the last thing you want to do is traverse the same 4 nodes
     and do four binary searches in each of those nodes for every read().

     For large fully cached data sets not caching B-Tree cursors will
     strip away 20-30% of your performance once your B-Tree depth exceeds
     4 or 5.  Also, the fan-out does not necessarily help there because
     the search within the B-Tree node costs almost as much as moving
     between B-Tree nodes.

     I found this out when I started comparing HAMMER performance with
     UFS.  For the fully cached case UFS was 30% faster until I started
     caching B-Tree cursors.  It was definitely noticable once my B-Tree
     grew past a million elements or so.  It disappeared completely when
     I started caching cursors into the B-Tree.

:>     If I were to do radix compression I would also want to go with a
:>     fully dynamic element size and fully dynamic fan-out in order to
:>     best-compress each B-Tree node.  Definitely food for thought.
:
:Indeed.  But Linux is braindamaged about large block size, so there is
:very strong motivation to stay within physical page size for the
:immediate future.  Perhaps if I get around to a certain hack that has
:been perenially delayed, that situation will improve:
:
:   "Variable sized page objects"
:   http://lwn.net/Articles/37795/

    I think it will be an issue for people trying to port HAMMER.  I'm trying
    to think of ways to deal with it.  Anyone doing an initial port can 
    just drop all the blocks down to 16K, removing the problem but
    increasing the overhead when working with large files.

:>     I'd love to do something like that.  I think radix compression would
:>     remove much of the topological bloat the B-Tree creates verses using
:>     blockmaps, generally speaking.
:
:Topological bloat?

    Bytes per record.  e.g. the cost of creating a small file in HAMMER
    is 3 B-Tree records (directory entry + inode record + one data record),
    plus the inode data, plus the file data.  For HAMMER that is 64*3 +
    128 + 112 (say the file is 100 bytes long, round up to a 16 byte
    boundary)... so that is 432 bytes.

    The bigger cost is when creating and managing a large file.  A 1 gigabyte
    file in HAMMER requires 1G/65536 = 16384 B-Tree elements, which comes
    to 1 megabyte of meta-data.  If I were to radix-compress those 
    elements the meta-data overhead would probably be cut to 300KB,
    possibly even less.

    Where this matters is that it directly effects the meta-data footprint 
    in the system caches which in turn directly effects the filesystem's
    ability to cache information without having to go to disk.  It can
    be a big deal.

:Ah, that is a very nice benefit of Tux3-style forward logging I forgot
:to mention in the original post: transaction size is limited only by
:available free space on the volume, because log commit blocks can be
:anywhere.  Last night I dreamed up a log commit block variant where the
:associated transaction data blocks can be physically discontiguous, to
:make this assertion come true, shortly after reading Stephen Tweedie's
:horror stories about what you have to do if you must work around not
:being able to put an entire, consistent transaction onto stable media
:in one go:

    Yes, that's an advantage, and one of the reasons why HAMMER is designed
    for large filesystems and not small ones.  Without a forward-log
    HAMMER has to reserve more media space for sequencing its flushes.

    Adding a forward log to HAMMER is possible, I might do it just for
    quick write()/fsync() style operations.  I am still very wary of the
    added code complexity.

:   http://olstrans.sourceforge.net/release/OLS2000-ext3/OLS2000-ext3.html
:
:Most of the nasties he mentions just vanish if:
:
:  a) File data is always committed atomically
:  b) All commit transactions are full transactions

    Yup, which you get for free with your forward-log.

:He did remind me (via time travel from year 2000) of some details I
:should write into the design explicitly, for example, logging orphan
:inodes that are unlinked while open, so they can be deleted on replay
:after a crash.  Another nice application of forward logging, which
:avoids the seek-happy linked list through the inode table that Ext3
:does.

    Orphan inodes in HAMMER will always be committed to disk with a
    0 link count.  The pruner deals with them after a crash.  Orphan
    inodes can also be commited due to directory entry dependancies.

:>     On the flip side, the reblocker doesn't just de-fragment the filesystem,
:>     it is also intended to be used for expanding and contracting partitions,
:>     and adding removing volumes in the next release.  Definitely a
:>     multi-purpose utility.
:
:Good point.  I expect Tux3 will eventually have a reblocker (aka
:defragger).  There are some really nice things you can do, like:
:
:  1) Set a new version so logical changes cease for the parent
:     version.
:
:  2) We want to bud off a given directory tree into its own volume,
:     so start by deleting the subtree in the current version.  If
:     any link counts in the directory tree remain nonzero in the
:     current version then there are hard links into the subtree, so
:     fail now and drop the new version.
:
:  3) Reblock a given region of the inode table tree and all the files
:     in it into one physically contiguous region of blocks
:
:  4) Add a free tree and other once-per volume structures to the new
:     home region.
:
:  5) The new region is now entirely self contained and even keeps its
:     version history.  At the volume manager level, map it to a new,
:     sparse volume that just has a superblock in the zeroth extent and
:     the new, mini filesystem at some higher logical address.  Remap
:     the region in the original volume to empty space and add the
:     empty space to the free tree.
:
:  6) Optionally reblock the newly budded filesystem to the base of the
:     new volume so utilties that do not not play well with sparse
:     volumes do not do silly things.

    Yes, I see.  The budding operations is very interesting to me...
    well, the ability to fork the filesystem and effectively write to
    a snapshot.  I'd love to be able to do that, it is a far superior
    mechanism to taking a snapshot and then performing a rollback later.

    Hmm.  I could definitely manage more then one B-Tree if I wanted to.
    That might be the ticket for HAMMER... use the existing snapshot
    mechanic as backing store and a separate B-Tree to hold all changes
    made to the snapshot, then do a merged lookup between the new B-Tree
    and the old B-Tree.  That would indeed work.

:>     So I'm not actually being cavalier about it, its just that I had to
:>     make some major decisions on the algorithm design and I decided to
:>     weight the design more towards performance and large-storage, and
:>     small-storage suffered a bit.
:
:Cavalier was a poor choice of words, the post was full of typos as well
:so I am not proud of it.  You are solving a somewhat different problem
:and you have code out now, which is a huge achievement.  Still I think
:you can iteratively improve your design using some of the techniques I
:have stumbled upon.  There are probably some nice tricks I can get from
:your code base too once I delve into it.

    Meh, you should see my documents when I post them without taking 10
    editorial passes.  Characters reversed, type-o's, sentences which make
    no sense, etc.

:>     In anycase, it sounds like Tux3 is using many similar ideas.  I think
:>     you are on the right track.
:
:Thankyou very much.  I think you are on the right track too, which you
:have a rather concrete way of proving.
:
:....
:>     rather then have to go through another 4 months of rewriting.  All
:>     major designs are an exercise in making trade-offs in topology, feature
:>     set, algorithmic complexity, debuggability, robustness, etc.  The list
:>     goes on forever.
:> 
:
:The big ahas! that eliminated much of the complexity in the Tux3 design
:were:
:
:  * Forward logging - just say no to incomplete transactions
:  * Version weaving - just say no to recursive copy on write
:
:Essentially I have been designing Tux3 for ten years now and working
:seriously on the simplifying elements for the last three years or so,
:either entirely on paper or in related work like ddsnap and LVM3.
:
:One of the nice things about moving on from design to implementation of
:Tux3 is that I can now background the LVM3 design process and see what
:Tux3 really wants from it.  I am determined to match every checkbox
:volume management feature of ZFS as efficiently or more efficiently,
:without violating the traditional layering between filesystem and
:block device, and without making LVM3 a Tux3-private invention.
:
:>     Laters!

    I can tell you've been thinking about Tux for a long time.  If I
    had one worry about your proposed implementation it would be in the
    area of algorithmic complexity.  You have to deal with the in-memory 
    cache, the log, the B-Tree, plus secondary indexing for snapshotted
    elements and a ton of special cases all over the place.  Your general
    lookup code is going to be very, very complex.

    My original design for HAMMER was a lot more complex (if you can
    believe it!) then the end result.  A good chunk of what I had to do
    going from concept to reality was deflate a lot of that complexity.
    When it got down to brass tacks I couldn't stick with using the 
    delete_tid as a B-Tree search key field, I had to use create_tid.  I
    couldn't use my fine-grained storage model concept because the
    performance was terrible (too many random writes interfering with
    streaming I/O).  I couldn't use a hybrid B-Tree, where a B-Tree element
    could hold the base of an entirely new B-Tree (it complicated the pruning
    and reblocking code so much that I just gave up trying to do it in
    frustration).  I couldn't implement file extents other then 16K and 64K
    blocks (it really complicated historical lookups and the buffer cache
    couldn't handle it) <--- that one really annoyed me.  I had overly
    optimized the storage model to try to get block pointers down to 32 bits
    by localizing B-Tree elements in the same 'super clusters' as the data
    they referenced.  It was a disaster and I ripped it out.  The list goes
    on :-)

    I do wish we had something like LVM on BSD systems.  You guys are very
    lucky in that regard.  LVM is really nice.

    BTW it took all day to write this!

					-Matt
					Matthew Dillon 
					<dillon@backplane.com>

:Hopefully not too much later.  I find this dialog very fruitful.  I just
:wish such dialog would occur more often at the design/development stage
:in Linux and other open source work instead of each group obsessively
:ignoring all "competing" designs and putting huge energy into chatting
:away about the numerous bugs that arise from rushing their design or
:ignoring the teachings of history.
:
:Regards,
:
:Daniel


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From: Daniel Phillips <phillips@...>
Subject: Re: [Tux3] Comparison to Hammer fs design
Date: Jul 27, 7:51 am 2008

linSubscribed now, everything should be OK.

On Friday 25 July 2008 19:02, Matthew Dillon wrote:
> :Yes, that is the main difference indeed, essentially "log everything" vs
> :"commit" style versioning.  The main similarity is the lifespan oriented
> :version control at the btree leaves.
> 
>     Reading this and a little more that you describe later let me make
>     sure I understand the forward-logging methodology you are using.
>     You would have multiple individually-tracked transactions in
>     progress due to parallelism in operations initiated by userland and each
>     would be considered committed when the forward-log logs the completion
>     of that particular operation?

Yes.  Writes tend to be highly parallel in Linux because they are
mainly driven by the VMM attempting to clean cache dirtied by active
writers, who generally do not wait for syncing.  So this will work
really well for buffered IO, which is most of what goes on in Linux.
I have not thought much about how well this works for O_SYNC or
O_DIRECT from a single process.  I might have to do it slightly
differently to avoid performance artifacts there, for example, guess
where the next few direct writes are going to land based on where the
most recent ones did and commit a block that says "the next few commit
blocks will be found here, and here, and here...".

When a forward commit block is actually written it contains a sequence
number and a hash of its transaction in order to know whether the
commit block write ever completed.  This introduces a risk that data
overwritten by the commit block might contain the same hash and same
sequence number in the same position, causing corruption on replay.
The chance of this happening is inversely related to the size of the
hash times the chance of colliding with the same sequence number in
random data times the chance of of rebooting randomly.  So the risk can
be set arbitrarily small by selecting the size of the hash, and using
a good hash.  (Incidentally, TEA was not very good when I tested it in
the course of developing dx_hack_hash for HTree.)

Note: I am well aware that a debate will ensue about whether there is
any such thing as "acceptable risk" in relying on a hash to know if a
commit has completed.  This occurred in the case of Graydon Hoare's
Monotone version control system and continues to this day, but the fact
is, the cool modern version control systems such as Git and Mercurial
now rely very successfully on such hashes.  Nonetheless, the debate
will keep going, possibly as FUD from parties who just plain want to
use some other filesystem for their own reasons.  To quell that
definitively I need a mount option that avoids all such commit risk,
perhaps by providing modest sized journal areas salted throughout the
volume whose sole purpose is to record log commit blocks, which then
are not forward.  Only slightly less efficient than forward logging
and better than journalling, which has to seek far away to the journal
and has to provide journal space for the biggest possible journal
transaction as opposed to the most commit blocks needed for the largest
possible VFS transaction (probably one).

>     If the forward log entries are not (all) cached in-memory that would mean
>     that accesses to the filesystem would have to be run against the log
>     first (scanning backwards), and then through to the B-Tree?  You
>     would solve the need for having an atomic commit ('flush groups' in
>     HAMMER), but it sounds like the algorithmic complexity would be
>     very high for accessing the log.

Actually, the btree node images are kept fully up to date in the page
cache which is the only way the high level filesystem code accesses
them.  They do not reflect exactly what is on the disk, but they do
reflect exactly what would be on disk if all the logs were fully
rolled up ("replay").

The only operations on forward logs are:

  1) Write them
  2) Roll them up into the target objects
  3) Wait on rollup completion

The rollup operation is also used for replay after a crash.

A forward log that carries the edits to some dirty cache block pins
that dirty block in memory and must be rolled up into a physical log
before the cache block can be flushed to disk.  Fortunately, such a
rollup requires only a predictable amount of memory: space to load
enough of the free tree to allocate space for the rollup log, enough
available cache blocks to probe the btrees involved, and a few
cache blocks to set up the physical log transaction.  It is the
responsibility of the transaction manager to ensure that sufficient
memory to complete the transaction is available before initiating it,
otherwise deadlock may occur in the block writeout path.  (This
requirement is the same as for any other transaction scheme).

One traditional nasty case that becomes really nice with logical
forward logging is truncate of a gigantic file.  We just need to
commit a logical update like ['resize', inum, 0] then the inode data
truncate can proceed as convenient.  Another is orphan inode handling
where an open file has been completely unlinked, in which case we
log the logical change ['free', inum] then proceed with the actual
delete when the file is closed or when the log is replayed after a
surprise reboot.

Logical log replay is not idempotent, so special care has to be taken
on replay to ensure that specified changes have not already been
applied to the target object.  I choose not to go the traditional route
of providing special case tests for "already applied" which get really
ugly or unreliable when there are lot of stacked changes.  Instead I
introduce the rule that a logical change can only be applied to a known
good version of the target object, which promise is fullfilled via the
physical logging layer.

For example, on replay, if we have some btree index on disk for which
some logical changes are outstanding then first we find the most recent
physically logged version of the index block and read it into cache,
then apply the logical changes to it there.  Where interdependencies
exist between updates, for example the free tree should be updated to
reflect a block freed by merging two btree nodes, the entire collection
of logical and physical changes has to be replayed in topologically
sorted order, the details of which I have not thought much about other
than to notice it is always possible.

When replay is completed, we have a number of dirty cache blocks which
are identical to the unflushed cache blocks at the time of a crash,
and we have not yet flushed any of those to disk.  (I suppose this gets
interesting and probably does need some paranoid flushing logic in
replay to handle the bizarre case where a user replays on a smaller
memory configuration than they crashed on.)  The thing is, replay
returns the filesystem to the logical state it was in when the crash
happened.  This is a detail that journalling filesystem authors tend
to overlook: actually flushing out the result of the replay is
pointless and only obscures the essential logic.  Think about a crash
during replay, what good has the flush done?

>     And even though you wouldn't have to group transactions into larger
>     commits the crash recovery code would still have to implement those
>     algorithms to resolve directory and file visibility issues.  The problem
>     with namespace visibility is that it is possible to create a virtually
>     unending chain of separate but inter-dependant transactions which either
>     all must go, or none of them.  e.g. creating a, a/b, a/b/c, a/b/x, a/b/c/d,
>     etc etc.

I do not see why this example cannot be logically logged in pieces:

   ['new', inum_a, mode etc] ['link', inum_parent, inum_a, "a"]
   ['new', inum_b, mode etc] ['link' inum_a, inum_b, "b"]
   ['new', inum_c, mode etc] ['link' inum_a_b, inum_c, "c"]
   ['new', inum_x, mode etc] ['link' inum_a_b, inum_x, "x"]
   ['new', inum_d, mode etc] ['link' inum_a_b_c, inum_d, "d"]

Logical updates on one line are in the same logical commit.  Logical
allocations of blocks to record the possibly split btree leaves and new
allocations omitted for clarity.  The omitted logical updates are
bounded by the depth of the btrees.  To keep things simple, the logical
log format should be such that it is impossible to overflow one commit
block with the updates required to represent a single vfs level
transaction.

I suspect there may be some terminology skew re the term "transaction".
Tux3 understands this as "VFS transaction", which does not include
trying to make an entire write (2) call atomic for example, but only
such things as allocate+write for a single page cache page or
allocate+link for a sys_link call.  Fsync(2) is not a transaction but
a barrier that Tux3 is free to realize via multiple VFS transactions,
which the Linux VFS now takes care of pretty well now after many years
of patching up the logic.

>     At some point you have to be able to commit so the whole mess 
>     does not get undone by a crash, and many completed mini-transactions
>     (file or directory creates) actually cannot be considered complete until
>     their governing parent directories (when creating) or children (when
>     deleting) have been committed.  The problem can become very complex.

Indeed.  I think I have identified a number of techniques for stripping
away much of that complexity.  For example, the governing parent update
can be considered complete as soon as the logical 'link' log commit has
completed.

>     Last question here:  Your are forward-logging high level operations.
>     You are also going to have to log meta-data (actual B-Tree manipulation)
>     commits in order to recover from a crash while making B-Tree
>     modifications.  Correct?

Correct!  That is why there are two levels of logging: logical and
physical.  The physical logging level takes care of updating the cache
images of disk blocks to match what the logical logging level expects
before it can apply its changes.  These two logging levels are
interleaved: where a logical change requires splitting a btree block,
the resulting blocks are logged logically, but linked into the parent
btree block using a logical update stored in the commit block of the
physical log transaction.  How cool is that?  The physical log
transactions are not just throwaway things, they are the actual new
data.  Only the commit block is discarded, which I suppose will leave
a lot of one block holes around the volume, but then I do not have to
require that the commit block be immediately adjacent to the body of
the transaction, which will allow me to get good value out of such
holes.  On modern rotating media, strictly linear transfers are not
that much more efficient than several discontiguous transfers that all
land relatively close to each other.

>     So your crash recovery code will have to handle 
>     both meta-data undo and completed and partially completed transactions.

Tux3 does not use undo logging, only redo, so a transaction is complete
as soon as there is enough information on durable media to replay the
redo records.

>     And there will have to be a tie-in between the meta-data commits and
>     the transactions so you know which ones have to be replayed.  That
>     sounds fairly hairy.  Have you figured out how you are doing to do that?

Yes.  Each new commit records the sequence number of the oldest commit
that should be replayed.  So the train lays down track in front of
itself and pulls it up again when the caboose passes.  For now, there
is just one linear sequence of commits, though I could see elaborating
that to support efficient clustering.

One messy detail: each forward log transaction is written into free
space wherever physically convenient, but we need to be sure that that
free space is not allocated for data until log rollup has proceeded
past that transaction.  One way to do this is to make a special check
against the list of log transactions in flight at the point where
extent allocation thinks it has discovered a suitable free block, which
is the way ddsnap currently implements the idea.  I am not sure whether
I am going to stick with that method for Tux3 or just update the disk
image of the free tree to include the log transaction blocks and
somehow avoid logging those particular free tree changes to disk.  Hmm,
a choice of two ugly but workable methods, but thankfully neither
affects the disk image.

> :Once I get down to the leaf level I binary search on logical address in
> :the case of a file index btree or on version in the case of an inode
> :table block, so this cost is still Log(N) with a small k.  For a
> :heavily versioned inode or file region this might sometimes result in
> :an overflow block or two that has to be linearly searched which is not
> :a big problem so long as it is a rare case, which it really ought to be
> :for the kinds of filesystem loads I have seen.  A common example of a
> :worst case is /var/log/messages, where the mtime and size are going to
> :change in pretty much every version, so if you have hourly snapshots
> :and hold them for three months it adds up to about 2200 16 byte inode
> :table attributes to record, about 8 inode table leaf blocks.  I really
> :do not see that as a problem.  If it goes up to 100 leaf blocks to
> :search then that could be a problem.
>     
>     I think you can get away with this as long as you don't have too many
>     snapshots, and even if you do I noticed with HAMMER that only a small
>     percentage of inodes have a large number of versions associated with
>     them from normal production operation.

Yes, that is my expectation.  I think everything will perform fine up
to a few hundred snapshots without special optimization, which is way
beyond current expectations.  Most users only recently upgraded to
journalling filesystems let alone having any exposure to snapshots
at all.

>     /var/log/messages 
>     is an excellent example of that.  Log files were effected the most though
>     I also noticed that very large files also wind up with multiple versions
>     of the inode, such as when writing out a terrabyte-sized file. 

Right, when writing the file takes longer than the snapshot interval.

>     Even with a direct bypass for data blocks (but not their meta-data,
>     clearly), HAMMER could only cache so much meta-data in memory before
>     it had to finalize the topology and flush the inode out.  A
>     terrabyte-sized file wound up with about 1000 copies of the inode
>     prior to pruning (one had to be written out about every gigabyte or so).

For Tux3 with an hourly snapshot schedule that will only be four or
five versions of the [mtime, size] attribute to reflect the 4.7 hours
write time or so, and just one version of each block pointer.

> :The penultimate inode table index block tells me how many blocks a
> :given inode lives in because several blocks will have the same inum
> :key.  So the lookup algorithm for a massively versioned file becomes:
> :1) read the first inode table block holding that inum; 2) read the last
> :block with the same inum.  The latter operation only needs to consult
> :the immediate parent index block, which is locked in the page cache at
> :that point.
> 
>     How are you dealing with expansion of the logical inode block(s) as
>     new versions are added?  I'm assuming you are intending to pack the
>     inodes on the media so e.g. a 128-byte inode would only take up
>     128 bytes of media space in the best case.  Multiple inodes would be
>     laid out next to each other logically (I assume), but since the physical
>     blocks are larger they would also have to be laid out next to each
>     other physically within any given backing block.  Now what happens
>     when one has to be expanded?

The inode table block is split at a boundary between inodes.

An "inode" is broken up into attribute groups arranged so that it makes
sense to update or version all the members of an attribute group
together.  The "standard" attribute group consists of mode, uid, gid,
ctime and version, which adds up to 16 bytes.  The smallest empty inode
(touch foo) is 40 bytes and a file with "foo" in it as immediate data
is 64 bytes.  This has a standard attribute, a link count attribute, a
size/mtime attribute, and an immediate data attribute, all versioned.
An inode with heavily versioned attributes might overflow into the next
btree leaf as described elsewhere.

Free inodes lying between active inodes in the same leaf use two bytes
each, a consequence of the inode leaf directory, which has a table
at the top of the leaf of two byte pointers giving the offset of each
inode in the leaf, stored backwards and growing down towards the inode
attributes.  (This is a slight evolution of the ddsnap leaf format.)

>     I'm sure this ties into the forward-log but even with the best algorithms
>     you are going to hit limited cases where you have to expand the inode.
>     Are you just copying the inode block(s) into a new physical allocation
>     then?

Split the inode in memory, allocating a new buffer; forward log the two
new pieces physically out to disk with a logical record in the commit
block recording where the two new pointers are to be inserted without
actually inserting them until a logical rollup episode is triggered.

> :>     I couldn't get away from having a delete_tid (the 'death version
> :>     numbers').  I really tried :-)  There are many cases where one is only
> :>     deleting, rather then overwriting.
> :
> :By far the most common case I would think.  But check out the versioned
> :pointer algorithms.  Surprisingly that just works, which is not exactly
> :obvious:
> :
> :   Versioned pointers: a new method of representing snapshots
> :   http://lwn.net/Articles/288896/
> 
>     Yes, it makes sense.  If the snapshot is explicitly taken then you
>     can store direct references and chain, and you wouldn't need a delete
>     id in that case.  From that article though the chain looks fairly
>     linear.  Historical access could wind up being rather costly.

I can think of three common cases of files that get a lot of historical
modifications:

  * Append log
      - The first iteration in each new version generates a new size
        attribute

  * Truncate/rewrite
     - The first iteration in each new version generates a new size
       attribute and either a new file index root or a new immediate
       data attribute

  * Database
     - The first file change in each new version generates a new
       versioned pointer at the corresponding logical address in the
       file btree and a new size attribute (just because mtime is
       bundled together with the size in one attribute group).

So the only real proliferation is the size/mtime attributes, which gets
back to what I was thinking about providing quick access for the
"current" version (whatever that means).

> :I was originally planning to keep all versions of a truncate/rewrite
> :file in the same file index, but recently I realized that that is dumb
> :because there will never be any file data shared by a successor version
> :in that case.  So the thing to do is just create an entirely new
> :versioned file data attribute for each rewrite, bulking up the inode
> :table entry a little but greatly constraining the search for versions
> :to delete and reducing cache pressure by not loading unrelated version
> :data when traversing a file.
> 
>     When truncating to 0 I would agree with your assessment.  For the 
>     topology you are using you would definitely want to use different
>     file data sets.  You also have to deal with truncations that are not
>     to 0, that might be to the middle of a file.  Certainly not a
>     common case, but still a case that has to be coded for.  If you treat
>     truncation to 0 as a special case you will be adding considerable
>     complexity to the algorithm.

Yes, the proposal is to treat truncation to exactly zero as a special
case.  It is a tiny amount of extra code for what is arguably the most
common file rewrite case.

>     With HAMMER I chose to keep everything in one B-Tree, whether historical
>     or current.  That way one set of algorithms handles both cases and code
>     complexity is greatly reduced.  It isn't optimal... large amounts of
>     built up history still slow things down (though in a bounded fashion).
>     In that regard creating a separate topology for snapshots is a good
>     idea.

Essentially I chose the same strategy, except that I have the file
trees descending from the inode table instead of stretching out to the
side.  I think this gives a more compact tree overall, and since I am
using just one generic set of btree operations to handle these two
variant btrees, additional code complexity is minimal.

> :>     Both numbers are then needed to 
> :>     be able to properly present a historical view of the filesystem.
> :>     For example, when one is deleting a directory entry a snapshot after
> :>     that deletion must show the directory entry gone.
> :
> :...which happens in Tux3 courtesy of the fact that the entire block
> :containing the dirent will have been versioned, with the new version
> :showing the entry gone.  Here is one of two places where I violate my
> :vow to avoid copying an entire block when only one data item in it
> :changes (the other being the atime table).  I rely on two things to
> :make this nice: 1) Most dirent changes will be logically logged and
> :only rolled up into the versioned file blocks when there are enough to
> :be reasonably sure that each changed directory block will be hit
> :numerous times in each rollup episode.  (Storing the directory blocks
> :in dirent-create order as in PHTree makes this very likely for mass
> :deletes.)  2) When we care about this is usually during a mass delete,
> :where most or all dirents in each directory file block are removed
> :before moving on to the next block.
> 
>     This could wind up being a sticky issue for your implementation.
>     I like the concept of using the forward-log but if you actually have
>     to do a version copy of the directory at all you will have to update the
>     link (or some sort of) count for all the related inodes to keep track
>     of inode visibility, and to determine when the inode can be freed and
>     its storage space recovered.

There is a versioned link count attribute in the inode.  When a link
attribute goes to zero for a particular version it is removed from the
inode.  When there are no more link attributes left, that means no
directory block references the inode for any version and the inode may
be reused.

Note: one slightly bizarre property of this scheme is that an inode
can be reused in any version in which its link count is zero, and the
data blocks referenced by different versions in the same file index
can be completely unrelated.  I doubt there is a use for that.

>     Directories in HAMMER are just B-Tree elements.  One element per
>     directory-entry.  There are no directory blocks.   You may want to
>     consider using a similar mechanism.  For one thing, it makes lookups
>     utterly trivial... the file name is hashed and a lookup is performed
>     based on the hash key, then B-Tree elements with the same hash key
>     are iterated until a match is found (usually the first element is the
>     match).  Also, when adding or removing directory entries only the
>     directory inode's mtime field needs to be updated.  It's size does not.

Ext2 dirents are 8 bytes + name + round up to 4 bytes, very tough to
beat that compactness.  We have learned through bitter experience that
anything other than an Ext2/UFS style physically stable block of
dirents makes it difficult to support NFS telldir cookies accurately
because NFS vs gives us only a 31 bit cookie to work with, and that is
not enough to store a cursor for, say, a hash order directory
traversal.  This is the main reason that I have decided to go back to
basics for the Tux3 directory format, PHTree, and make it physically
stable.

In the PHTree directory format lookups are also trivial: the directory
btree is keyed by a hash of the name, then each dirent block (typically
one) that has a name with that hash is searched linearly.  Dirent block
pointer/hash pairs are at the btree leaves.  A one million entry
directory has about 5,000 dirent blocks referenced by about 1000 btree
leaf blocks, in turn referenced by three btree index blocks (branching
factor of 511 and 75% fullness).  These blocks all tend to end up in
the page cache for the directory file, so searching seldom references
the file index btree.

I did some back of the envelope calculations of the number of cache
lines that have to be hit for a lookup by Hammer with its fat btree
keys and lower branching factor, vs Tux3 with its high branching
factor, small keys, small pointers, extra btree level for dirent
offset stability and stupidly wasteful linear dirent search.  I will
not bore the list with the details, but it is obvious that Tux3/PHTree
will pass Hammer in lookup speed for some size of directory because of
the higher btree fanout.  PHTree starts from way behind courtesy of the
32 cache lines that have to be hit on average for the linear search,
amounting to more than half the CPU cost of performing a lookup in a
million element directory, so the crossover point is somewhere up in
the millions of entries.

Thus annoyed, I cast about for a better dirent leaf format than the
traditional Ext2 directory block, and found one after not too much
head scratching.  I reuse the same leaf directory format as for inode
leaf blocks, but this time the internal offsets are sorted by lexical
name order instead of inode number order.  The dirents become Ext2
dirents minus the record number format, and minus the padding to
four byte alignment, which does not do anything useful.  Dirent inum
increases to 6 bytes, balanced by saving 1.5 pad bytes on average,
so the resulting structure stores about 2% fewer dirents than the
Ext2 format against a probable 50% reduction in CPU latency per lookup.
A fine trade indeed.

The new format is physically stable and suitable for binary searching.
When we need to manage space within the leaf for creating or deleting
an entry, a version of the leaf directory ordered by offset can be
created rapidly from the lexically sorted directory, which occupies at
most about six cache lines.  The resulting small structure can be
cached to take advantage of the fact that most mass creates and deletes
keep hitting the same dirent block repeatedly, because Tux3 hands the
dirents to getdents in physical dirent order.

I concluded that both Hammer and PHTree are exceptionally fast at name
resolution.  When I have the new dirent block format in place I expect
to really hammer Hammer.  But then I do not expect Hammer to stand
still either ;-)

>     Your current directory block method could also represent a problem for
>     your directory inodes... adding and removing directory entries causing
>     size expansion or contraction could require rolling new versions
>     of the directory inode to update the size field.  You can't hold too
>     much in the forward-log without some seriously indexing.  Another
>     case to consider along with terrabyte-sized files and log files.

A PHTree directory grows like an append-only file, a block at a time,
though every time a new entry is added, mtime has to change, so mtime
changes more often than size.  They are grouped together in the same
attribute, so the distinction is moot.  If the directory is modified at
least once after every snapshot (quite likely) then the snapshot
retention policy governs how many mtime/size attributes are stored in
the inode.

Say the limit is 1,024 snapshots, then 16K worth of those attributes
have to be stored, which once again encourages me to store the "current"
mtime specially where it can be retrieved quickly.  It would also be a
good idea to store the file data attribute in front of the mtime/size
attributes so only stat has to worry about the bulky version info, not
lookup.

Note that there is no such thing as revving an entire inode, only bits
and pieces of it.

For truly massive number of versions, the plan is to split up the
version tree into subtrees, each having a thousand versions or so.  For
each version subtree (hmm, subversion tree...) there is a separate
inode btree carrying only attributes and file data owned by members of
that subtree.  At most log(subtrees) of those tables need to be
accessed by any versioned entity algorithm.  (Note the terminology
shift from versioned pointers to versioned entities to reflect the fact
that the same algorithms work for both pointers and attributes.)  I
think this approach scales roughly to infinity, or at least to the
point where log(versions) goes vertical which is essentially never.  It
requires a bounded amount of implementation effort, which will probably
be deferred for months or years.

Incidentally, PHTree directories are pretty much expand-only, which
is required in order to maintain physical dirent stability.  Compaction
is not hard, but it has to be an admin-triggered operation so it will
not collide with traversing the directory.

> :>     Both numbers are 
> :>     also needed to be able to properly prune out unwanted historical data
> :>     from the filesystem.  HAMMER's pruning algorithm (cleaning out old  
> :>     historical data which is no longer desired) creates holes in the
> :>     sequence so once you start pruning out unwanted historical data
> :>     the delete_tid of a prior record will not match the create_tid of the
> :>     following one (historically speaking).
> :
> :Again, check out the versioned pointer algorithms.  You can tell what
> :can be pruned just by consulting the version tree and the create_tids
> :(birth versions) for a particular logical address.  Maybe the hang is
> :that you do not organize the btrees by logical address (or inum in the
> :case of the inode table tree).  I thought you did but have not read
> :closely enough to be sure.
> 
>     Yes, I see.  Can you explain the versioned pointer algorithm a bit more,
>     it looks almost like a linear chain (say when someone is doing a daily
>     snapshot).  It looks great for optimal access to the HEAD but it doesn't
>     look very optimal if you want to dive into an old snapshot. 

It is fine for diving into old snapshots.  Lookup is always O(elements)
where an element is a versioned pointer or (later) extent, which are
very compact at 8 bytes each.  Because the truncate/rewrite file update
case is so common, you will rarely see more than one version at any
given logical address in a file.  A directory growing very slowly could
end up with about 200 different versions of each of its final few
blocks, one for each added entry.  It takes on the order of a
microsecond to scan through a few hundred versioned pointers to find
the one that points at the version in which we are interested, and that
only happens the first time the directory block is read into the page
cache.  After that, each reference to the block costs a couple of
lookups in a page cache radix tree.

The version lookup algorithm is not completely obvious: we scan through
the list of pointers looking for the version label that is nearest the
version being accessed on the path from that version to the root of the
version tree.  This potentially expensive search is accelerated using a
bitmap table to know when a version is "on the path" and a pre-computed
ord value for each version, to know which of the versions present and
"on the path" for a given logical address is furthest from the root
version.  This notion of furthest from the root on the path from a given
version to the root implements data inheritance, which is what yields
the compact representation of versioned data, and is also what
eliminates the need to explicitly represent the death version.  You have
discovered this too, in a simpler form.  I believe that once you get the
aha you will want to add versions of versions to your model.

>     For informational purposes: HAMMER has one B-Tree, organized using
>     a strict key comparison.  The key is made up of several fields which
>     are compared in priority order:
> 
> 	localization	- used to localize certain data types together and
> 			  to separate pseudo filesystems created within
> 			  the filesystem.
> 	obj_id		- the object id the record is associated with.
> 			  Basically the inode number the record is 
> 			  associated with.
> 
> 	rec_type	- the type of record, e.g. INODE, DATA, SYMLINK-DATA,
> 			  DIRECTORY-ENTRY, etc.
> 
> 	key		- e.g. file offset
> 
> 	create_tid	- the creation transaction id
> 
>     Inodes are grouped together using the localization field so if you
>     have a million inodes and are just stat()ing files, the stat
>     information is localized relative to other inodes and doesn't have to
>     skip file contents or data, resulting in highly localized accesses
>     on the storage media.
> 
>     Beyond that the B-Tree is organized by inode number and file offset.
>     In the case of a directory inode, the 'offset' is the hash key, so
>     directory entries are organized by hash key (part of the hash key is
>     an iterator to deal with namespace collisions).
> 
>     The structure seems to work well for both large and small files, for
>     ls -lR (stat()ing everything in sight) style traversals as well as 
>     tar-like traversals where the file contents for each file is read.

That is really sweet, provided that you are ok with the big keys.  It
was smart to go with that regular structure, giving you tons of control
over everything and get the system up and running early, thus beating the
competition.  I think you can iterate from there to compress things and
be even faster.  Though if I were to presume to choose your priorities
for you, it would be to make the reblocking optional.

>     The create_tid is the creation transaction id.  Historical accesses are
>     always 'ASOF' a particular TID, and will access the highest create_tid
>     that is still <= the ASOF TID.  The 'current' version of the filesystem
>     uses an asof TID of 0xFFFFFFFFFFFFFFFF and hence accesses the highest
>     create_tid.

So you have to search through a range of TIDs to find the operative one
for a particular version, just as I have to do with versioned pointers.
Now just throw in a bitmap like I use to take the version inheritance
into account and you have versioned pointers, which means snapshots of
snapshots and other nice things.  You're welcome :-)

I think that versioned pointer lookup can be implemented in O(log(E))
where E is the number of versioned pointers (entities) at the same
logical address, which would put it on a par with btree indexing, but
more compact.  I have sketched out an algorithm for this but I have not
yet tried to implement it.  

>     There is also a delete_tid which is used to filter out elements that
>     were deleted prior to the ASOF TID.  There is currently an issue with
>     HAMMER where the fact that the delete_tid is *not* part of the B-Tree
>     compare can lead to iterations for strictly deleted elements, verses
>     replaced elements which will have a new element with a higher create_tid
>     that the B-Tree can seek to directly.  In fact, at one point I tried
>     indexing on delete_tid instead of create_tid, but created one hell of a
>     mess in that I had to physically *move* B-Tree elements being deleted
>     instead of simply marking them as deleted.

This is where I do not quite follow you.  File contents are never
really deleted in subsequent versions, they are just replaced.  To
truncate a file, just add or update size attribute for the target
version and recover any data blocks past the truncate point that
belongs only to the target version.

Extended attributes and dirents can be truly deleted.  To delete an
extended attribute that is shared by some other version I will insert
a new versioned attribute that says "this attribute is not here" for
the current version, which will be inherited by its child versions.

Dirent deletion is handled by updating the dirent block, possibly
versioning the entire block.  Should the entire block ever be deleted
by a directory repacking operation (never actually happens on the vast
majority of Linux systems) that will be handled as a file truncate.

I see that in Hammer, the name is deleted from the btree instead, so
fair enough, that is actual deletion of an object.  But it could be
handled alternatively by adding a "this object is not here" object,
which might be enough to get rid of delete_tid entirely and just have
create_tid.

> :Fair enough.  I have an entirely different approach to what you call
> :mirroring and what I call delta replication.  (I reserve the term
> :mirroring to mean mindless duplication of physical media writes.)  This
> :method proved out well in ddsnap:
> :
> :   http://phunq.net/ddtree?p=zumastor/.git;a=tree;h=fc5cb496fff10a2b03034fcf95122f5828149...
> :
> :(Sorry about the massive URL, you can blame Linus for that;-)
> :
> :What I do in ddsnap is compute all the blocks that differ between two
> :versions and apply those to a remote volume already holding the first
> :of the two versions, yielding a replica of the second version that is
> :logically but not physically identical.  The same idea works for a
> :versioned filesystem: compute all the leaf data that differs between
> :two versions, per inum, and apply the resulting delta to the
> :corresponding inums in the remote replica.  The main difference vs a
> :ddsnap volume delta is that not all of the changes are physical blocks,
> :there are also changed inode attributes, so the delta stream format
> :has to be elaborated accordingly.
> 
>     Are you scanning the entire B-Tree to locate the differences?  It
>     sounds you would have to as a fall-back, but that you could use the
>     forward-log to quickly locate differences if the first version is
>     fairly recent.

I plan to scan the whole btree initially, which is what we do in ddsnap
and it works fine.  But by looking at the mtime attributes in the inode
I can see in which versions the file data was changed and thus not have
to scan most files.  Eventually, building in some accelerator to be able
to skip most inode leaves as well would be a nice thing to do.  I will
think about that in the background.

>     HAMMER's mirroring basically works like this:  The master has synchronized
>     up to transaction id C, the slave has synchronized up to transaction id A.
>     The mirroring code does an optimized scan of the B-Tree to supply all
>     B-Tree elements that have been modified between transaction A and C.
>     Any number of mirroring targets can be synchronized to varying degrees,
>     and can be out of date by varying amounts (even years, frankly).
> 
>     I decided to use the B-Tree to optimize the scan.  The B-Tree is
>     scanned and any element with either a creation or deletion transaction
>     id >= the slave's last synchronization point is then serialized and
>     piped to the slave.

That is nearly identical to the ddsnap replication algorithm, and we
also send the list over a pipe.  Ddsnap does not deal with attributes,
only volume blocks, otherwise I think we arrived at the same thing.

>     To avoid having to scan the entire B-Tree I perform an optimization
>     whereby the highest transaction id laid down at a leaf is propagated
>     up the B-Tree all the way to the root.  This also occurs if a B-Tree
>     node is physically deleted (due to pruning), even if no elements are
>     actually present at the leaf within the transaction range.
>     Thus the mirroring scan is able to skip any internal node (and its
>     entire sub-tree) which has not been modified after the synchronization
>     point, and is able to identify any leaf for which real, physical deletions
>     have occured (in addition to logical deletions which simply set the
>     delete_tid field in the B-Tree element) and pass along the key range
>     and any remaining elements in that leaf for the target to do a
>     comparative scan with.

Hmm.  I have quite a few bits available in the inode table btree node
pointers, which are currently only going to be free inode density
hints.  Something to think about.

>     This allows incremental mirroring, multiple slaves, and also allows
>     a mirror to go offline for months and then pop back online again
>     and optimally pick up where it left off.  The incremental mirroring
>     is important, the last thing I want to do is have to scan 2 billion
>     B-Tree elements to do an incremental mirroring batch.

Very, very nice.

>     The advantage is all the cool features I got by doing things that way.
>     The disadvantage is that the highest transaction id must be propagated
>     up the tree (though it isn't quite that bad because in HAMMER an entire
>     flush group uses the same transaction id, so we aren't constantly
>     repropagating new transaction id's up the same B-Tree nodes when
>     flushing a particular flush group).

This is a job for forward logging :-)

>     You may want to consider something similar.  I think using the
>     forward-log to optimize incremental mirroring operations is also fine
>     as long as you are willing to take the 'hit' of having to scan (though
>     not have to transfer) the entire B-Tree if the mirror is too far
>     out of date.

There is a slight skew in your perception of the function of forward
logging here, I think.  Forward logs transactions are not long-lived
disk objects, they just serve to batch together lots of little updates
into a few full block "physical" updates that can be written to the
disk in seek-optimized order.  It still works well for this application.
In short, we propagate dirty bits up the btree while updating the btree
disk blocks only rarely.  Instead, the node dirty bits are tucked into
the commit block of the write transaction or namespace edit that caused
the change, and are in turn propagated into the commit block of an
upcoming physical rollup, eventually working their way up the on-disk
btree image.  But they are present in the cached btree image as soon as
they are set, which is the structure consulted by the replication
process.

> :I intend to log insertions and deletions logically, which keeps each
> :down to a few bytes until a btree rollup episode comes along to perform
> :updating of the btree nodes in bulk.  I am pretty sure this will work
> :for you as well, and you might want to check out that forward logging
> :trick.
> 
>     Yes.  The reason I don't is because while it is really easy to lay down
>     a forward-log, intergrating it into lookup operations (if you don't
>     keep all the references cached in-memory) is really