On 09/12/2010 10:18 PM, Anthony Liguori wrote:
But since you have to boot before you can run any serious test, if
it takes 5 seconds to do an fsck(), it's highly likely that it's
not even noticeable.
What if it takes 300 seconds?
That means for a 1TB disk you're taking 500ms per L2 entry, you're
fully allocated and yet still doing an fsck. That seems awfully
unlikely.
I meant for a fully populated L1. That's 10ms per L2.
Your math is off. A single L2 entry covers 2GB worth of logical
space. That means a 1TB image consists of 512 L2s. 300 / 512 == .585
which is 585ms.
That's a fully populated L1 on a 1TB image.
I meant fully populated L1, all 32K entries. Not the 1TB disk.
But since that's 64TB, that's unlikely too. It can still take 10s
for a 2TB disk.
Ah, you're talking about a 64TB image. Recall that we can read L2s in
parallel. I have trouble imaging that we'd get serialized performance
with a 64TB backing store. It's much more likely you've got more than
one spindle in this scenario.
When you have more than one spindle, you also have more than one guest.
If they're all fscking at the same time, it actually gets slower.
Why is in n^2? It's still n*m. If your image is 4TB instead of
100GB, the time increases by a factor of 40 for both.
It's n*m but either n ~= m in which case it's n^2 or m << n, in
which case, it's just n, or m >> n in which case, it's just O(m).
This is where asymptotic complexity ends up not being terribly
helpful :-)
Let me put this another way though, if you support internal
snapshots, what's a reasonable number of snapshots to expect
reasonable performance with? 10? 100? 1000? 10000?
I'd say 10. Not that I really want to support internal snapshots, it
doesn't work well with multiple disks.
I don't think that's reasonable. The folks that I've talked to about
snapshots seem to want to do crazy things like use it for
checkpointing. TBH, I think they're looking for the ability to do
thousands of checkpoints with an efficient way to release old
checkpoints.
I imagine that's the design point things like btrfs are trying to
achieve.
Well, qcow2 isn't going to get anywhere close to that. For that many
you basically have to have extents and back references, or your metadata
explodes.
On the other hand, linear L2 (which now become L1) means your fsck is
just a linear scan of the table, which is probably faster than qcow2
allocation...
And this is just a data layout optimization which is the sort of thing
that we should let performance data drive.
Well, if you drop L1, the code becomes simpler.
For a 1PB disk image with qcow2, the reference count table is
128GB. For a 1TB image, the reference count table is 128MB. For a
128GB image, the reference table is 16MB which is why we get away
with it today.
Anytime you grow the freelist with qcow2, you have to write a brand
new freelist table and update the metadata synchronously to point to
a new version of it. That means for a 1TB image, you're potentially
writing out 128MB of data just to allocate a new cluster.
s/freelist/refcount table/ to translate to current qcow2
nomenclature. This is certainly not fast. You can add a bunch of
free blocks each time you mitigate the growth but I can't of many
circumstances where a 128MB write isn't going to be noticeable. And
it only gets worse as time moves on because 1TB disk images are
already in use today.
That's a strong point. qcow2 doubles on each allocation, it
amortizes, but the delay is certainly going to be noticable.
You can do it ahead of time (so guest writes don't need to wait) but
it's still expensive.
The trouble is, safe growth of the reference count table is hard
because it's contiguous. That means you need to copy the table to
another location all at once instead of just creating a new L1 table
and reusing most of the existing L2 entries.
You can use the same technique as streaming. You initiate a copy in the
background. Updates to the region already copied go to both copies,
updates to the region not yet copied go to the original copy.
It's a damning flaw in the format for large images. You can
preallocate the whole thing up front to try to avoid the cost at run
time but even then, that's a huge cost to pay in disk space up front.
I'm more or less convinced it qcow2 is unusable for large images without
incompatible format changes.
It's very easy to neglect the details in something like qcow2.
We've been talking like the refcount table is basically free to read
and write but it's absolutely not. With large disk images, you're
caching an awful lot of metadata to read the refcount table in fully.
If you reduce the reference count table to exactly two bits, you can
store that within the L1/L2 metadata since we have an extra 12 bits
worth of storage space. Since you need the L1/L2 metadata anyway,
we might as well just use that space as the authoritative source of
the free list information.
The only difference between qcow2 and qed is that since we use an
on-demand table for L1/L2, our free list may be non-contiguous.
Since we store virtual -> physical instead of physical->virtual, you
have to do a full transversal with QED whereas with qcow2 you may
get lucky. However, the fact that the reference count table is
contiguous in qcow2 is a design flaw IMHO because it makes growth
extremely painful with large images to the point where I'll claim
that qcow2 is probably unusable by design with > 1TB disk images.
If you grow it in the background, it should be usable; since it
happens once every 1TB worth of writes, it's not such a huge load.
I'll agree this is increasing complexity.
Trouble is, the reference count table is your authoritative source of
whether something is free. You can grow it in the background but if
you need to allocate clusters before your done growing, you have to
stall the request. Those stalls can get very long on large disk images.
No, you update both copies. It's hardly fun though.
We can optimize qed by having a contiguous freelist mapping
physical->virtual (that's just a bitmap, and therefore considerably
smaller) but making the freelist not authoritative. That makes it
much faster because we don't add another sync and let's us fallback
to the L1/L2 table for authoritative information if we had an
unclean shutdown.
It's a good compromise for performance and it validates the qed
philosophy. By starting with a correct and performant approach that
scales to large disk images, we can add features (like unmap)
without sacrificing either.
How would you implement the bitmap as a compatible feature?
First, a refresher on the consistency model.
Metadata is not sync'd to disk when initially written because we try
to avoid having stronger metadata consistency than data consistency.
We are forced to sync to enforce ordering when updating L1 with a new
L2 but that's rare. If a guest attempts to sync data, we sync
metadata too.
Because we may have unsync'd metadata, all L1 and L2 entries need to
be scanned to try to find unreachable entries based on the start up
file size *before* allocating any new clusters upon start-up (noting
that we can read and rewrite existing clusters).
This is a correct design and performant this is where we start from.
The start up cost is undesirable, so to reduce the need to scan all
entries up front, we add a feature that introduces a meta-data clean
flag in the header. The meta-data clean flag tells an implementation
that there were no cluster allocations since the last sync() which
means cluster allocation can now happen without searching for and
correcting unreachable entries.
I don't think you can add it as a compatible feature. Implementations
without the feature will load a clean image and dirty it without
clearing the flag. So it has to be in from day 1.
An implementation can now set the meta-data clean flag right before
any sync() operation
What happens if the clean flag reaches the disk before a metadata write
(or truncate) makes it? That's a false-clean which can lead us to avoid
a scan which would be needed.
and unset the flag before the first cluster allocation (being careful
to sync the unsetting of the flag). This eliminates a lot of
unnecessary scans in the case of safe shut down.
Why do you need to sync clearing the flag? Worst case it is lost and
you have a false-dirty, leading to an unnecessary scan.
We can also set a timer after unsetting the meta-data clean flag to
sync() after, say, 5 minutes. This further decreases the number of
times we have to scan so that the fsck() window only happens when
power failure occurs within 5 minutes of the last cluster allocation.
N.B., this flag is purely an optional feature that is backwards
compatible. If an implementation ignores the meta-data clean flag, it
just always scans for unreachable entries at startup.
No, you can get a false-clean by an old qemu crashing.
I think this sort of stuff can be done using a mount count. Every mount
increments the mount count. "clean" sets a clean counter equal to the
mount count. "dirty" sets the clean counter to 0. You could do this
for any number of recoverable metadata structures (like a dirty bitmap).
This is still a correct design and we've eliminated the vast majority
of start up scans.
The freelist would also be an optional feature. If the freelist
pointer is non-zero in the header, it means that an implementation can
find free blocks by reading the location of the freelist pointer in
the header.
The same issue - old qemu mounts, updates disk but ignores freelist, new
qemu mounts and uses stale freelist.
A mount count fixes this (an alternative to a mount count is a new flags
field with 1 bit per optional feature, any unrecognized compatible
feature is cleared in the flags field, letting a new implementation
detect that an old qemu saw and ignored the feature).
We maintain the freelist in memory and we write it out right before
any sync() operation. Adding to the freelist in memory (i.e. UNMAP)
would clear the freelist pointer on disk. We could use this as an
opportunity to schedule a future sync() just like as above.
We would actually write free list changes to disk while the freelist
pointer was set to 0 such that when we did a sync(), we were just
marking the freelist valid again.
An implementation that doesn't know about free list has to do a full
scan to determine free blocks. That said, an implementation can also
just ignore free blocks entirely and always allocate from the end.
The freelist would be a cluster that also happens to be an free'd
entry. It would then contain a list of free clusters and would look
something like:
struct freelist
{
uint64_t next_cluster; /* zero to represent EOL */
uint64_t num_entries;
uint64_t entry[num_entries]; /* cluster offsets of free blocks */
};
A bitmap would work better vs fragmentation - the worst case for this
structure is pretty bad. For 64k clusters, each 1TB needs just 2MB of
space, which is easy to scan and even reasonable to hold in memory.
It's important to not represent the full list in a contiguous region
for the reasons discussed in other notes re: refcount table.
Background copying can work around the issue, but yes.
This remains a completely correct implementation. We add no
additional syncs above what we already have and we write out the
freelist changes on demand. The only window that would require a
rebuild of the freelist would be a hard power failure 5 minutes since
the last unmap.
After the last allocation; unmap only results in a leak (which can be
recovered in the background).
One gotcha is that the freelist is double metadata which means that
freeing a block requires an ordered write with respect to the L2 table
update. Otherwise, you could allocate something off of the freelist
that an L2 entry still pointed to.
So there's a sync() in the UNMAP path, but that's unavoidable without
a lot of cleverness. You could potentially delay the sync until you
reallocate the block but it's not clear that UNMAP needs to be fast
(yet).
I don't see why you need the sync. If the freelist update is lost,
you've leaked a block which you can recover with a rebuild. If the L2
update is lost, you have the freelist pointing into a used cluster, but
you can recover using fsck. If both are lost, nothing happened.
--
error compiling committee.c: too many arguments to function
