On Tue, Jul 4, 2017 at 1:39 PM, Xavier Hernandez <xhernandez@xxxxxxxxxx> wrote:
Hi Pranith,
On 03/07/17 05:35, Pranith Kumar Karampuri wrote:
Ashish, Xavi,
I think it is better to implement this change as a separate
read-after-write caching xlator which we can load between EC and client
xlator. That way EC will not get a lot more functionality than necessary
and may be this xlator can be used somewhere else in the stack if possible.
while this seems a good way to separate functionalities, it has a big problem. If we add a caching xlator between ec and *all* of its subvolumes, it will only be able to cache encoded data. So, when ec needs the "cached" data, it will need to issue a request to each of its subvolumes and compute the decoded data before being able to use it, so we don't avoid the decoding overhead.
Also, if we want to make the xlator generic, it will probably cache a lot more data than ec really needs. Increasing memory footprint considerably for no real use.
Additionally, this new xlator will need to guarantee that the cached data is current, so it will need its own locking logic (that would be another copy&paste of the existing logic in one of the current xlators) which is slow and difficult to maintain, or it will need to intercept and reuse locking calls from parent xlators, which can be quite complex since we have multiple xlator levels where locks can be taken, not only ec.
This is a relatively simple change to make inside ec, but a very complex change (IMO) if we want to do it as a stand-alone xlator and be generic enough to be reused and work safely in other places of the stack.
If we want to separate functionalities I think we should create a new concept of xlator which is transversal to the "traditional" xlator stack.
Current xlators are linear in the sense that each one operates only at one place (it can be moved by reconfiguration, but once instantiated, it always work at the same place) and passes data to the next one.
A transversal xlator (or maybe a service xlator would be better) would be one not bound to any place of the stack, but could be used by all other xlators to implement some service, like caching, multithreading, locking, ... these are features that many xlators need but cannot use easily (nor efficiently) if they are implicitly implemented in some specific place of the stack outside its control.
The transaction framework we already talked, could be though as one of these service xlators. Multithreading could also benefit of this approach because xlators would have more control about what things can be processed by a background thread and which ones not. Probably there are other features that could benefit from this approach.
In the case of brick multiplexing, if some xlators are removed from each stack and loaded as global services, most probably the memory footprint will be lower and the resource usage more optimized.
I like the service xlator approach. But I don't think we have enough time to make it operational in the short term. Let us go with implementation of this feature in EC for now. I didn't realize the extra cost of decoding when I thought about the separation. So I guess we will stick to the old idea for now.
Just an idea...
Xavi
On Fri, Jun 16, 2017 at 4:19 PM, Ashish Pandey <aspandey@xxxxxxxxxx
<mailto:aspandey@xxxxxxxxxx>> wrote:
I think it should be done as we have agreement on basic design.
------------------------------------------------------------ ------------
*From: *"Pranith Kumar Karampuri" <pkarampu@xxxxxxxxxx
<mailto:pkarampu@xxxxxxxxxx>>
*To: *"Xavier Hernandez" <xhernandez@xxxxxxxxxx
<mailto:xhernandez@xxxxxxxxxx>>
*Cc: *"Ashish Pandey" <aspandey@xxxxxxxxxx
<mailto:aspandey@xxxxxxxxxx>>, "Gluster Devel"
<gluster-devel@xxxxxxxxxxx <mailto:gluster-devel@gluster.org >>
*Sent: *Friday, June 16, 2017 3:50:09 PM
*Subject: *Re: Disperse volume : Sequential Writes
On Fri, Jun 16, 2017 at 3:12 PM, Xavier Hernandez
<xhernandez@xxxxxxxxxx <mailto:xhernandez@xxxxxxxxxx>> wrote: <mailto:xhernandez@xxxxxxxxxx
On 16/06/17 10:51, Pranith Kumar Karampuri wrote:
On Fri, Jun 16, 2017 at 12:02 PM, Xavier Hernandez
<xhernandez@xxxxxxxxxx <mailto:xhernandez@xxxxxxxxxx>
<mailto:xhernandez@xxxxxxxxxx>>> wrote:
On 15/06/17 11:50, Pranith Kumar Karampuri wrote:
On Thu, Jun 15, 2017 at 11:51 AM, Ashish Pandey
<aspandey@xxxxxxxxxx <mailto:aspandey@xxxxxxxxxx>
<mailto:aspandey@xxxxxxxxxx <mailto:aspandey@xxxxxxxxxx>>
<mailto:aspandey@xxxxxxxxxx
<mailto:aspandey@xxxxxxxxxx> <mailto:aspandey@xxxxxxxxxx
<mailto:aspandey@xxxxxxxxxx>>>> wrote:
Hi All,
We have been facing some issues in disperse (EC)
volume.
We know that currently EC is not good for random
IO as it
requires
READ-MODIFY-WRITE fop
cycle if an offset and offset+length falls in
the middle of
strip size.
Unfortunately, it could also happen with
sequential writes.
Consider an EC volume with configuration 4+2.
The stripe
size for
this would be 512 * 4 = 2048. That is, 2048
bytes of user data
stored in one stripe.
Let's say 2048 + 512 = 2560 bytes are already
written on this
volume. 512 Bytes would be in second stripe.
Now, if there are sequential writes with offset
2560 and of
size 1
Byte, we have to read the whole stripe, encode
it with 1
Byte and
then again have to write it back.
Next, write with offset 2561 and size of 1 Byte
will again
READ-MODIFY-WRITE the whole stripe. This is
causing bad
performance.
There are some tools and scenario's where such
kind of load is
coming and users are not aware of that.
Example: fio and zip
Solution:
One possible solution to deal with this issue is
to keep
last stripe
in memory.
This way, we need not to read it again and we
can save READ fop
going over the network.
Considering the above example, we have to keep
last 2048 bytes
(maximum) in memory per file. This should not
be a big
deal as we already keep some data like xattr's
and size info in
memory and based on that we take decisions.
Please provide your thoughts on this and also if
you have
any other
solution.
Just adding more details.
The stripe will be in memory only when lock on the
inode is active.
I think that's ok.
One
thing we are yet to decide on is: do we want to read
the stripe
everytime we get the lock or just after an extending
write is
performed.
I am thinking keeping the stripe in memory just after an
extending write
is better as it doesn't involve extra network operation.
I wouldn't read the last stripe unconditionally every
time we lock
the inode. There's no benefit at all on random writes
(in fact it's
worse) and a sequential write will issue the read anyway
when
needed. The only difference is a small delay for the
first operation
after a lock.
Yes, perfect.
What I would do is to keep the last stripe of every
write (we can
consider to do it per fd), even if it's not the last
stripe of the
file (to also optimize sequential rewrites).
Ah! good point. But if we remember it per fd, one fd's
cached data can
be over-written by another fd on the disk so we need to also
do cache
invalidation.
We only cache data if we have the inodelk, so all related fd's
must be from the same client, and we'll control all its writes
so cache invalidation in this case is pretty easy.
There exists the possibility to have two fd's from the same
client writing to the same region. To control this we would need
some range checking in the writes, but all this is local, so
it's easy to control it.
Anyway, this is probably not a common case, so we could start by
caching only the last stripe of the last write, ignoring the fd.
May be implementation should consider this possibility.
Yet to think about how to do this. But it is a good point.
We should
consider this.
Maybe we could keep a list of cached stripes sorted by offset in
the inode (if the maximum number of entries is small, we could
keep the list not sorted). Each fd should store the offset of
the last write. Cached stripes should have a ref counter just to
account for the case that two fd's point to the same offset.
When a new write arrives, we check the offset stored in the fd
and see if it corresponds to a sequential write. If so, we look
at the inode list to find the cached stripe, otherwise we can
release the cached stripe.
We can limit the number of cached entries and release the least
recently used when we reach some maximum.
Yeah, this works :-).
Ashish,
Can all of this be implemented by 3.12?
One thing I've observed is that a 'dd' with block size
of 1MB gets
split into multiple 128KB blocks that are sent in
parallel and not
necessarily processed in the sequential order. This
means that big
block sizes won't benefit much from this optimization
since they
will be seen as partially non-sequential writes. Anyway
the change
won't hurt.
In this case as per the solution we won't cache anything
right? Because
we didn't request anything from the disk. We will only keep
the data in
cache if it is not aligned write which is at the current
EOF. At least
that is what I had in mind.
Suppose we are writing multiple 1MB blocks at offset 1. If each
write is split into 8 blocks of 128KB, all writes will be not
aligned, and can be received in any order. Suppose that the
first write happens to be at offset 128K + 1. We don't have
anything cached, so we read the needed stripes and cache the
last one. Now the next write is at offset 1. In this case we
won't get any benefit from the previous write, since the stripe
we need is not cached. However the write from the user point of
view is sequential.
It won't hurt but it won't take all benefits from the new
caching mechanism.
As a mitigating factor, we could consider to extend the previous
solution I've explained to allow caching multiple stripes per
fd. A small number like 8 would be enough.
Xavi
Xavi
---
Ashish
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--
Pranith
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Pranith
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Pranith
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Pranith
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