On 03/11/2013 01:22 PM, Ivan Shapovalov wrote:
On 02 March 2013 17:55:48 Edward Shishkin wrote:
On 02/23/2013 01:21 PM, Ivan Shapovalov wrote:
[...]
But here's what I currently think about discard implementation.
In filesystems like jfs, it is implemented pretty straightforward.
"Online" discard on block freeing is done through hooking into
function dbFree(), which marks the blocks as free in the _working_
allocation map. Batch discard via FITRIM ioctl is done through locking
the whole allocation group, allocating everything in it, trimming
these blocks and freeing them again.
For reiser4, I think it will translate into something like this:
With "online" discard, it would be better to discard the blocks at
transaction commit time (the time when working bitmap is copied to the
persistent one... am I right?)
I am sorry, but I still don't know the TRIM/discard background well
enough to make any decisions. I understand that a file system should
issue some commands to "help" the hardware? What those commands will
result in?
---- tl;dr area begin
The TRIM is a command in the ATA protocol, operating on a sector range.
It tells the hardware (storage) that the given sector range is not used
anymore and hence data contained in it can be discarded/removed.
(Similar commands exist in several other protocols, like SCSI UNMAP and
SD ERASE, and the "discard" is an in-kernel abstraction to all such
commands.)
The reason why do we need such a command for SSDs is that in flash memory
an "overwrite data" operation is actually an "erase + write data" and is
much more costly than just a "write data onto free space". Flash memory
is organized into pages (usually 4K), which are further grouped into
blocks (512K); and while a write is done per-page, an erase is done
per-block (so a controller shall read the whole block into cache and then
rewrite all pages in it, except the one being updated).
Modern controllers do internal block remapping to achieve some "wear
leveling" (i. e. spreading use across all blocks instead of continuously
rewriting one block which is updated by the user), but they obviously
need a pool of free blocks, and anyway - writes to the locations that the
software would consider empty still may trigger a read-erase-write cycle.
So, the TRIM command notifies the controller that the block can be erased
and returned to the free pool. There is a restriction on sector ranges
given to the command: they should actually represent whole blocks
(otherwise they are ignored, AFAIK).
Hello Ivan.
Thanks for the background. This is exactly what did I want to see.
So, from the software's point of view, an SSD-aware operation looks like
1) putting whatever is likely to be updated simultaneously into the same
block (TRIM unit);
Not sure if I understand the (1). Could you please say more?
I wanted to say that we could tune block allocation algorithms (or whatever is
responsible for choosing a specific free block from all available ones) so
that
the data which is likely to be updated together, e. g. file body and stat-
data,
will be placed in blocks of the same erase unit (== TRIM unit, as reported by
kernel). It will just make less read-erase-write cycles when the buffers are
written.
But well, this is no more than just a heuristic. The kernel sometimes just
fails to provide a sane granularity. (Again, in my case, granularity is
reported as 1 sector/512 bytes.)
Ah, I understand. Good remark.
Stat-data can be logically prepended to file body via
REISER4_3_5_KEY_ALLOCATION policy. Also we need to make
sure that stat-data and file body are in the same extent of physical
blocks. The single existing block allocation policy would be OK
for this, but I afraid that we'll need to fix it a bit for SSD (to not
"spoil" discard extents, see below).
2) delaying writeback in hope that more adjacent data will be written at
once;
Yes. In reiser4 we delay everything what is possible.
And, I think, discard requests shouldn't be an exception..
3) notify the storage when the blocks are logically freed by issuing a
TRIM
command.
(1) and (2) are largely my guesses (and anyway out of scope), while
(3) is a common practice and is implemented at storage driver, kernel and
filesystem layers.
---- tl;dr area end
So, inside the filesystem we need to notify the kernel about we need to
implement TRIM (more precisely, discard - as we're working with
in-kernel abstractions) support in the filesystem
About the implementation:
There is an API call, blkdev_issue_discard() [1], which does all the
work and is supposed to be called from the filesystem. The discard
properties are stored in struct queue_limits.
And for the filesystem itself, there are generally two modes to support
discard operations [2].
1) "Realtime" or online discard - the filesystem discards blocks as they
are deallocated (files being deleted, tree nodes being cut, etc.).
There is another source of deallocated blocks in reiser4, that you
should be aware of. This is the flush procedure. This procedure
operates on a reiser4 atom and is called every time before its commit
to complete all delayed actions:
(1) allocate all extents of the atom (for files manages by
unix-file plugin);
(2) compress data of the atom (for files managed by
cryptcompress file plugin);
(3) balance tree in the atom's locality;
(4) schedule commit policy for dirty blocks of the atom
(relocate, or overwrite).
(3) - (4) are sources of deallocated blocks: (3) will release blocks
freed after squeezing an atom. And (4) will be the most active issuer
of discard requests: at this phase we determine the best allocation
for the whole group of atom's dirty blocks in accordance with some
heuristic. And it can happen that a lot of blocks will change their
on-disk locations (they will be assigned to so-called atom's "relocate
set"). Other dirty blocks (which won't change their on-disk locations)
are assigned to atom's "overwrite set".
Committing an atom in reiser4 looks like this:
(a) write atom's relocate set (simply write the blocks to their new
locations on disk);
(b) write atom's overwrite set (via journal, aka "wandering logs"),
i.e. at first, write the dirty blocks to journal, then overwrite the
blocks at their old locations on disk;
(c) update system records to indicate, that transaction is
completed.
And thank _you_ for the reiser4 background. Now its "pipeline" is more clear
to me...
This is really good news. Nevertheless, I think you probably have
troubles in understanding "who is who". Note, that in the real life we
don't wait for flush completion to perform (a). We write the relocate
set gradually as we proceed with flush (see write_flush_queue()).
This allows to free some occupied pages quickly in the situations
of memory pressure. But freeing pages of the overwrite set is a
looong story..
I think that in "realtime" mode we should issue all discard requests
of an atom at the point after (b) and before (c). Indeed, at this point
all updated bitmaps are successfully committed, so in the worst case
(power off when issuing a series of discard requests) we'll just loose
only a part of discard requests (not fatal).
Yes, seems good. BTW - if we implement realtime discard in such way, will we
automatically get discard at transaction replay?
You are looking at the root of things, but I don't promise automatic
discard at journal replay ;)
To perform discard at journal replay, we need to encode
information about extents (which should be discarded) to the
journal. In the meanwhile, reiser4 journal is a rather stupid thing,
just a sequence of blocks with journal header and journal footer.
The question is: do we really need to discard anything at journal
replay? What will happen if we won't discard some extents in a
situation (crash, power loss, etc), which is rare enough?
I think that nothing criminal. After all, user can run the real-time
discard to make sure that everything is discarded properly.
Or is there no such thing as
"transaction replay" in r4?
I would prefer to say "journal replay", meaning that it completes
a transaction. Transaction is a complex thing, which includes
relocate and overwrite sets. Journal contains only overwrite set.
2) "Batch" discard - the filesystem discards all free blocks upon a user's
request (when mounted).
In this "batch" case, the signaling is done through a FITRIM ioctl on any
file.
"Batch" mode:
Implementing it should be simple enough (if I'm making correct assumptions
about how does reiser4 work): we can just lock the bitmap and walk through
it, issuing a discard for each long enough free sequence.
Mmm, I haven't found definition of "free block"..
For example, we have deleted a file by unlink(2), and the transaction,
which contains the updated bitmap is not yet committed. And here is
an interesting question: at this moment blocks of that file are free, or
busy? ;)
Does reiser4 have a notion of "effective" bitmap? The one which
represents current on-disk data,
Yes.
Reiser4 maintains 2 in-memory copies of a modified
bitmap block: before and after modifications.
without any in-flight transactions.
I've been thinking of this:
- lock transactions from being committed
- get the "effective" bitmap
- directly scan it and issue discard requests
IMHO it would be incorrect. If a bitmap block has been
modified, then why should we look at the old (effective) copy?
Let's discard the final result of things..?
- unlock everything
I can show you a list of reiser4 locks. This is scaring.
Let's avoid locks, especially giant ones, whenever it is possible..
Is it possible? If not, then actually the algorithm you described in a
follow-up message (separate process) looks viable and optimal.
BTW, that is partially similar to how other filesystems implement batch
discard - they use existing interfaces to (temporarily) allocate blocks
in a loop and then discard these allocated blocks.
"Realtime" mode:
It will be more complex given that we have to do the actual work on
transaction commit.
You are right about the slowness of bitmap comparison (yes, 32K bitops...
I
haven't thought about it); we'll need to store locations to discard in
some
per-atom data structure.
Let's define a "minimal discard range" to be a block range,
1) whose begin is properly aligned,
2) whose size is equal to discard granularity.
This can be checked using data from struct queue_limits (exact algorithm
can be derived from code of blkdev_issue_discard()).
Actually, simply storing each deallocated interval in the atom and then
iterating through the list upon commit will be suboptimal.
Reasons:
- if a single deallocated range is smaller than the discard granularity,
then this particular range won't be discarded even if it is surrounded by
enough free blocks to make a minimal discard range;
- we won't be able to merge small adjacent ranges to form a range that's
long enough.
Solution:
- record all deallocated ranges verbatim (in a list);
- on commit time, for each recorded range find minimal discard range(s)
which encompass the given range and check if all their blocks can be
discarded (i. e. are free);
- add each suitable minimal discard range to a locally-allocated tree
(while merging the added ranges);
Why not to just maintain per-atom rb-trees? All deallocated ranges
will be represented as records (extents) in those trees. It looks more
simple, no?
When truncate(2) deallocates a range of blocks, we find a position in
such "discard tree", and try to merge this range with neighbouring
extents. If they are not mergeable, then insert one more extent...
I see the following (hope resolvable) problems here:
1. Ranges of blocks freed by truncate(2) can be "spoiled" by
relocate decisions performed in flush time (action (4) above).
I mean the situation when the flush procedure borrows block
numbers for the "best allocation" from... our discard extents.
In other words, before issuing a discard request, we need to
check our discard extents for possible "holes". Such check can
be also implemented by the updated bitmap, which is contained
in the same atom.
2. Another problem is maintenance of the "discard trees" during
atom's evolution. Sometimes atoms may merge. So their "discard
trees" should be respectively merged. For the beginning we can
merge trees for by simply allocating a new empty one and placing
there all extents from the trees we want to merge (N+M operatioins).
Later we can implement "rb-trees with fingers", invented for fast
merge, which will take only log(min{N,M}) operations [1].
3. And one more problem: it would be better to not allocate anything
at flush and commit time: usually flush/commit is a reiser4 respond
to memory pressure notifications of the operating system. Linux
doesn't have any reservation mechanisms for subsystems, which
need memory to free memory.
At flush time we'll need memory to represent deallocated ranges as
records in the "discard trees". I think it makes sense to preallocate
special per-atom pools for those needs. I think 20-40K per atom should
be enough.
Here's the problem I was going to resolve with my algorithm:
- discard granularity is e. g. 4 blocks
- Initial bitmap: 1001 (1 - busy, 0 - free)
- deallocate first and last blocks, not necessarily in a single transaction
- resulting bitmap: 0000 (discard ranges: 0:1 and 3:1)
- ranges can't be merged since they aren't adjacent (even if in the same
transaction) and can't be discarded since they all are smaller than the
granularity, while the whole 4-block range can be discarded easily.
Right before discard we can perform the following operation
on our rb-tree:
for each extent:
check its "locality" of discard-unit-size by [1];
replace it with the largest possible extent;
try to merge with the previous "extended" extent;
Now we can discard the resulted tree with a clear conscience.
Another solution:
every time when inserting an extent when deallocating blocks
(by ->truncate(), etc), check the "locality" of the extent by [1],
and replace it by the largest possible extent. In this case no
additional passes are needed before discard, but, I think that
the first way is more preferable (seven troubles - one response).
However, in the second way we can save some memory, though...
[1] There is a number of fast bit operations, see e.g. nlz, ntz:
http://en.wikipedia.org/wiki/Find_first_set
Also, by delaying "range -> TRIM unit" conversion to commit time
we get solution of (1) for granted since we already access bitmap - and such
accesses also aren't going to be very expensive.
TBH, i don't see how we get solution of (1) for granted.
The flush procedure reallocates blocks of atom (assigns new block numbers
for them).
It means that flush must operate with our discard extents:
a) add new extents (when releasing old blocks);
b) split existing extents (when assigning new blocks);
We can minimize (b), or even avoid it completely by tweaking the current
block
allocation policy (say, try to find new block numbers beyond all
"active" atoms).
However (a) will be tons and tons of operations: we'll use COW transaction
model for SSD, so that every dirty block of data will always get new
location
on disk.
So in atoms we can maintain simple linked lists, and (2) is solved too because
they can be merged for constant time.
BTW, how are we going to handle allocation of new blocks in the
scheme with a simple linked list of deallocated extents?
Moreover, I've seen filesystems adding an artificial granularity limit
(i. e. do not even bother to discard ranges smaller than N sectors) to aid
performance in case if kernel-reported limits are wrong.
In our case, we can also do that without sacrificing long-term efficiency
(as all freed ranges will be eventually discarded once they become long
enough).
How does that sound?
Regarding memory - I wonder if multiple atoms can be flushed concurrently.
Yes.
Moreover, one atom can be flushed by many flushers :)
We can specify number of flushers per atom (1 by default).
If no, then preallocated pools for per-atom lists + global (per-mount) pool
for the resulting discard tree.
BTW, maybe we can infer the per-atom pool size from atom_max_size?
Good idea.
But let's start with identical pools for simplicity..
- issue discard for all found ranges.
Hope this won't be too slow. BTW, kernel sometimes seems to report wrong
granularity. In my case, granularity is reported as 512 bytes.
So we can make a recap.
Batched discard:
Some clarifications are needed to understand if we can implement
something useful here..
Realtime discard:
Now It is more or less clear, how to implement it in reiser4. You will
want to make a friendship with reiser4 transaction manager. This is
rather advanced and complicated thing (with this manager reiser4
has much more capabilities, than any other file system). Start with
understanding, that every cached block (page) of reiser4 partition is
contained in some atom: this is captured by reiser4 transaction
manager (try_capture() and friends). Note, that atom contains not
only dirty blocks. Clean blocks also participate in relations created by
transaction manager (see [2] for details). Once in a while (responding
on memory pressure notification, or because the transaction is too
large/old) atoms get committed: their subsets of dirty blocks are
written to disk by steps (a, b, c) above.
You will encounter specific problems, but experience shows all they
are resolvable.
Hope they are... So I see following steps:
- Access the atom from bitmap manipulation plugin
- Store freed ranges to the in-atom tree/list
- Traverse through the transaction manager and add code supporting the discard
lists (merge, etc - if any)
- Patch the flush procedure to perform discard requests after writing blocks
Am I missing something?
I would do things in the following order:
1) fully define data structures for discard extents (assume we'll call
it "dt" (discard tree));
add a respective pointer to the struct atom;
2) define and encode operations on those data structures, i.e.:
->dt_init();
->dt_alloc_extent(start, number_of_blocks); // what happens if an
extent of blocks is allocated;
->dt_free_extent(start, number_of_blocks); // what happens if an
extent of blocks is freed;
->dt_release();
etc.
3) find places in the reiser4 code where:
(a) ->truncate() deallocates blocks of pruned files;
(b) flush deallocates blocks freed after squeezing an atom;
(c) flush reallocates blocks when making decision about best
allocation;
4) call respective hooks (2) at the found places (a, b, c). Note, that
at those
places we have a pointer to the atom.
5) discard the extents in the atom commit time and release the atom's dt.
Let me know, if you need a help with any items above.
Thanks,
Edward.
[1]http://citeseerx.ist.psu.edu/viewdoc/summary?doi=10.1.1.38.4454
[2]http://lwn.net/2001/1108/a/reiser4-transaction.php3
[...]
by performing a comparison between the
old (on-disk) and new bitmaps, remembering all changed chunks and
issuing discard for them.
I afraid that comparison the bitmaps is something expensive: it means
4K*8 = 32K comparisons per bitmap block.. Maybe it makes sense to
accumulate the "difference" in special per-atom data structures
(say, rb-trees)?
Also, the discard granularity can be higher
than the bitmap granularity. E. g. if we have a bitmap pattern like
"0010" and it changes to "0000", it would be better to issue a discard
for 4 blocks instead of just one.
And with FITRIM, we could just lock the bitmap and walk through it,
discarding all free chunks. Of course, it can only be done if locking
policy allows us to "just lock the bitmap"...
BTW, I'm afraid I don't understand what "a proposal" means. Is it a
kind of some official document - and if yes, who needs it?
Nothing official, this is a usual practice in groups that work
remotely: someone send a kind of roadmap. In the simplest case it
can be a set of links where one can read about TRIM/discard.
Maybe "proposal" sounds too official? :)
For the other things: the freezing issue seems to be related to
fsync() indeed; the freeze rate decreased substantially when I stopped
using InnoDB as the MySQL backend. Some of them remained, seemingly
related to Dropbox (== concurrent reads and writes to the same file).
This is a known problem, I'll try to find Reiser's suggestions how to
resolve this..
Due to transactional fs's nature?
And yes, I'll try to do the bisection as soon as enough free time
appears... Will a virtual machine be enough, or it is crucial that the
tests shall be performed on a real machine?
It can be remote, but it should be a real machine. BTW, where are you
territorially?
I'm in Moscow (RU). Actually, I can do that on my primary PC - if those
old
kernels are able to boot a SandyBridge chipset.
BTW, mirror at mirror.sit.wisc.edu is offline... I'll use
mirror.linux.org.au - and hope that patches will apply to any of the
intermediate states. What is the first known bad version?
Ivan.
Edward.
Thanks,
Ivan.
2013/2/10 Edward Shishkin<edward.shishkin@xxxxxxxxx>:
Hi Ivan,
How our TRIM/dsicard is doing?
Any questions, or everything is clear? :)
Edward.
On 01/17/2013 05:39 PM, Edward Shishkin wrote:
On 01/07/2013 02:42 AM, Ivan Shapovalov wrote:
Hi again Edward,
Hello.
Here's what I want to try to do with reiser4 in meantime. I'd
appreciate some
hints on that all...
So, first thing I'd like to implement is TRIM/discard support, both
online
(via -o discard) and in a separate FITRIM ioctl().
That's just because I've got an SSD two days ago and thus now have to
use in
rootfs some discard-aware fs like ext4.
I think it would be nice for beginning. Moreover, reiser4 still
doesn't
have any setup optimal for SSD.
Unfortunately I don't have a ready proposal for TRIM/discard support
in
reiser4.
I have ready proposals for the following features (they can be rather
complicated for the beginners though):
1) Repacker (On-line defragmentation);
2) Support of different transaction models:
a. pure journalling;
b. pure COW (Copy-On-Write);
c. smart (the current "mixed" one);
d. no transaction support (for people with UPSs);
3) Subvolumes (AKA "chunkfs");
4) Snapshots.
And then I want to do something with performance: sometimes during
heavy I/O
to a slow /home storage (especially when it's multithreaded) many
processes,
including the DE, just get stuck in "D" state and sit there for a
minute or
two with load average of apx. 5.5 (on a hyperthreaded 2-core CPU).
and some process waits for fsync() completion?
For the first, I can look into other filesystems' implementations,
but
I'll
probably be unsure at which layer to put the actual discard call (in
order not
to break reiser4's transactional nature).
If you decide to proceed with TRIM/discard support, you will need to
prepare the proposal by yourself. Let's start with some background,
that is:
. clarify underlying reasons (specific for SSD geometry?) of
TRIM/discard support: why do we need such support on the file
system layer;
. review of existing hardware and software means for such support;
. etc..
And yes, it would be nice to review existing TRIM/discard support
implementations in other file systems (say, ext4).
Once we figure out, what bits of reiser4 you should understand
perfectly to implement TRIM/discard support, I'll provide you with
respective hints.
And for the second, I just don't know why does that happen. Can it be
due to
some r4-specific things/issues or that's just a horribly slow random
access
speed of my hw?
Which hw? SSD?
I also remember complaints that umount (i.e. the final sync takes 2-3,
or even more minutes). It looks like in some cases reiser4 accumulates
too much dirty stuff..
It would be nice to periodically dump some info about atoms (current
number of all atoms, size of each atom, etc) to see the full picture
of
their evolution during such freezing. I think, it makes sense to port
the old reiser4 profiling stuff, and populate it with more info (if
needed).
Also there is an oldest issue:
The following (old) benchmarks created with mongo(*) test suit show x2
advantage of reiser4 against reiserfs(v3) on CREATE phase (let's
consider only this phase for simplicity):
http://web.archive.org/web/20061113154648/http://www.namesys.com/bench
ma
rks.html
I've made similar benchmarks with latest 2.6 kernels (sorry, lost the
results) and found that the advantage has disappeared (real time in
CREATE phase is the same as of reiserfs, or even worse). It shouldn't
be so: it indicates that something wrong is going on.. I remember
people complained on the performance drop in reiser4 long time ago,
but
didn't have a chance to investigate this.
The straightforward way to narrow down the problem changeset is to
bisect starting from 2.6.8-mm2, the archives can be found here:
http://mirror.sit.wisc.edu/pub/linux/kernel/people/akpm/patches/2.6/
http://ftp.icm.edu.pl/packages/linux-reiserfs/reiser4-for-2.6/
http://mirror.sit.wisc.edu/pub/linux/kernel/people/edward/reiser4/reis
er
4-for-2.6/
However, it can be rather painful and requires a separate machine.
Thanks,
Edward.
(*)
http://sourceforge.net/projects/reiser4/files/reiser4-utils/bench-stre
ss
-tools/
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