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Unformatted text preview: USENIX Association Proceedings of the
FAST 2002 Conference on
File and Storage Technologies
Monterey, California, USA
January 28-30, 2002 THE ADVANCED COMPUTING SYSTEMS ASSOCIATION © 2002 by The USENIX Association
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This copyright notice must be included in the reproduced paper. USENIX acknowledges all trademarks herein. Safety, Visibility, and Performance in a Wide-Area File System
Minkyong Kim, Landon P. Cox, and Brian D. Noble
Electrical Engineering and Computer Science
University of Michigan
minkyong,lpcox,bnoble @umich.edu Abstract
As mobile clients travel, their costs to reach home ﬁling services change, with serious performance implications. Current
ﬁle systems mask these performance problems by reducing
the safety of updates, their visibility, or both. This is the result of combining the propagation and notiﬁcation of updates
from clients to servers.
Fluid Replication separates these mechanisms. Client updates
are shipped to nearby replicas, called WayStations, rather than
remote servers, providing inexpensive safety. WayStations and
servers periodically exchange knowledge of updates through
reconciliation, providing a tight bound on the time until updates are visible. Reconciliation is non-blocking, and update
contents are not propagated immediately; propagation is deferred to take advantage of the low incidence of sharing in ﬁle
Our measurements of a Fluid Replication prototype show that
update performance is completely independent of wide-area
networking costs, at the expense of increased sharing costs.
This places the costs of sharing on those who require it, preserving common case performance. Furthermore, the beneﬁts
of independent update outweigh the costs of sharing for a workload with substantial sharing. A trace-based simulation shows
that a modest reconciliation interval of 15 seconds can eliminate 98% of all stale accesses. Furthermore, our traced clients
could collectively expect availability of ﬁve nines, even with
deferred propagation of updates. 1 Introduction
Mobile devices have become an indispensable part of
the computing infrastructure. However, networking costs
continue to render them second class citizens in a distributed ﬁle system. Limits imposed by networking costs
are not new, but they are no longer due to “last mile”
constraints. With the widespread deployment of broadband connectivity, mobile users often ﬁnd themselves in
a neighborhood of good network performance. Unfortunately, they must reach back across the wide area to
interact with their ﬁle servers; the latency and congestion along such paths impose a substantial performance
penalty. To cope with increased costs, wide-area and mobile ﬁle
systems employ two techniques to limit their use of the
remote server. The ﬁrst technique is caching. When
workloads have good locality and infrequent sharing,
caching can avoid most ﬁle fetches from the server.
The second technique is optimistic concurrency control.
Clients defer shipping updated ﬁles until time and resources permit. Ideally, most updates will be overwritten
at the client, and need never be sent to the server.
Deferring updates improves performance, but harms
safety and visibility. An update is safe if it survives the
theft, loss, or destruction of the mobile client that created
it. Safety requires that the contents of an update reside
on at least one other host. An update is visible if every
other client in the system knows it exists. Visibility requires that notiﬁcation of an update reaches all replicas
in the system.
Current ﬁle systems needlessly combine safety and visibility by propagating the contents of an update, implicitly notifying the destination that the update exists.
Fluid Replication separates the concerns of safety and
visibility through the addition of secondary replica sites.
These sites, called WayStations, act as server replicas
to nearby clients, servicing uncached reads and writes.
Client writes are propagated to a WayStation immediately, providing safety. WayStations periodically reconcile their updates with the servers for which they act
as replicas; this exchanges update notiﬁcations, but not
the contents of those updates. Since reconciliation involves only meta-data, it can be done frequently, providing bounded visibility. Put another way, Fluid Replication aggressively writes updates back to the WayStation,
but periodically invalidates updates between WayStation
To maintain a simple consistency model, Fluid Replication provides copy semantics; each replica behaves as if a
reconciliation actually copied all new versions accepted
by the other replica. To support this, each WayStation retains ﬁle versions in escrow in the event that they are referenced at other replicas. Escrow also allows us to provide wait-free reconciliations with no additional mechanism. We have built a prototype of Fluid Replication, and evaluated it with a variety of benchmarks. Updates are isolated from wide-area networking costs in the absence of
sharing; clients pay wide-area costs only for accesses to
shared ﬁles. A trace-based simulation of Fluid Replication shows that a modest reconciliation interval of ﬁfteen seconds provides a stale access rate of only 0.01%,
compared to a sharing rate of 0.6%. Space devoted to escrow storage is modest but bursty, with a high water mark
of less than 10 MB. The impact of deferred propagation
on availability is also small; our traced client population
would expect to see one failed access per year, collectively. 2 Related Work
A number of projects have explored ﬁle systems for mobile, wide-area clients. Many of the ideas in Fluid Replication stem from the Coda ﬁle system. Coda provides
high availability through disconnected operation 
for clients and server replication  between wellconnected servers. Fluid Replication is orthogonal to
both; we do not expect clients to avoid disconnection or
server failure. Coda also supports weakly-connected operation  to utilize low bandwidth between a client and
its servers. During weakly-connected operation, Coda
clients defer updates, and ship them later through a process called trickle reintegration. This sends both the
knowledge of an update and its contents to the server.
Since this is expensive, Coda defers reintegration in the
hopes that some updates will be canceled by overwrites
or deletions. This aging window is set to ten minutes to
capture an acceptable fraction of possible optimizations,
trading network bandwidth for more aggressive propagation.
Ficus [9, 23] shares Coda’s goal of providing optimistic
ﬁle access, but uses a peer-to-peer architecture. Each update is accepted by a single replica and asynchronously
propagated to other sites, but no effort is made to ensure
that update messages arrive. To propagate missed updates, Ficus provides copy semantics through the actual
exchange of updates between two replicas. Updates are
discovered by a disk scan at each replica, a heavyweight
process. Reconciliation exchanges both knowledge and
content of updates. Since this is expensive, it is intended
only for well-connected peers, making visibility dependent on the mobility pattern of clients.
Bayou  provides optimistic concurrency control for
database applications, where distributed updates are
eventually committed by a primary replica. Deno 
extends the Bayou model to provide distributed commit.
These systems exchange full update sets whenever convenient. Unlike Ficus, they log updates; unlike Coda, updates are represented by-operation rather than by-value. This inherently combines notiﬁcation of an update and
its contents. However, since databases exhibit many ﬁnegrained, partial updates, this may be less of an issue than
for ﬁle systems. Exchanges between peers rely on client
mobility and communication patterns. So, these systems
can offer eventual visibility, but cannot bound the time
required to do so.
OceanStore  advocates an architecture for global
storage. Like Fluid Replication, it envisions storage supported by a loose confederation of independent servers.
A primary focus of OceanStore is safely using untrusted
nodes in the infrastructure. This is also an important
problem for Fluid Replication, which we address elsewhere . OceanStore’s consistency mechanism, like
that of Ficus and Bayou, is based on epidemic algorithms, and so cannot bound visibility. OceanStore represents updates by operation rather than by value, combining notiﬁcation of updates with the shipment of their
TACT  provides a middleware service to coordinate the consistency of data used by wide-area applications. Unlike Fluid Replication, it combines visibility and safety of updates. However, it does provide
strong bounds on visibility. In addition to temporal
bounds—which are offered by Fluid Replication—it can
also bound the number of unseen writes or the weight of
those writes. Combining Fluid Replication’s separation
of safety and visibility with TACT’s mechanisms to trigger reconciliation can offer applications a tunable consistency mechanism with isolation from wide-area networking costs.
xFS is a ﬁle system designed for tightly coupled clusters of workstations . However, early designs focused
on the wide area . In this model, clients were aggregated based on physical proximity, and served by a
consistency server. On update, a client would retain the
contents of the update—reducing safety—and notify the
consistency server that the update occurred. Consistency
between these second-level replicas and the main storage server was to be managed conservatively, imposing
wide-area networking costs on normal operation.
Finally, JetFile  is motivated by the same concerns
as the early xFS design, though it provides a number
of additional features. It uses a central version server
that provides serialization, but does not hold data; objects are stored primarily on the client that creates them,
reducing safety. Clients locate ﬁles with scalable reliable
multicast . As in Ficus, invalidations are sent besteffort. However, in JetFile, the central version server
broadcasts invalidations periodically to ensure bounded
consistency; this period provides consistency equivalent
to Fluid Replication. JetFile’s most serious drawback is
that it depends on ubiquitous IP multicast. 3 Design
This section presents the design of Fluid Replication with
an eye towards separating the concerns of safety and visibility. Clients treat WayStations, wide-area replica sites,
exactly as they would a server. Each WayStation reconciles its updates with the server periodically; reconciliation exchanges notiﬁcation of updates, but not their
contents. In order to provide copy semantics without actually copying during reconciliation, replicas hold reconciled updates in escrow. Due to the escrow mechanism,
reconciliation does not require any node to hold a lock
during a network round trip. 3.1 Context Fluid Replication is designed to complement a clientserver architecture, such as that of AFS . In the
AFS model, servers provide the centralized point of administration. They are expected to be carefully maintained to provide the best possible availability to their
clients. Clients consist of workstations and mobile devices. Clients are considered the property of their users;
they are not carefully administered and are much more
likely to fail. This is particularly true of mobile clients.
The AFS model uses callbacks for consistency. Before
using an uncached ﬁle, a client must ﬁrst fetch it from
the server. As a side effect of the fetch, the server establishes a callback on that ﬁle. The client can then use the
now-cached copy. If a client modiﬁes a ﬁle, the new version is sent to the server on close; this is called a store
event. The server breaks callback on any other clients
caching this ﬁle, invalidating their copies, before accepting the store. Unix semantics forbids invalidation of an
open ﬁle; such ﬁles are invalidated after they are closed.
If a ﬁle subject to pending invalidation is modiﬁed, the
close returns an error. 3.2 WayStations: Hosting Replicas The success of Fluid Replication depends on WayStations, which act as replicas for servers. They provide
caching and immediate safety to nearby clients without
penalizing performance, and act as reconciliation partners with remote servers. A WayStation is able to host
replicas of any service for any client, regardless of administrative domain, and can host replicas of multiple
In our architecture, WayStations play a role similar to
that of servers; one model for a WayStation is as a feebased service, much as broadband connectivity is for
travelers today. In contrast to clients, they are carefully administered machines, and are relatively permanent members of the infrastructure. While they can be
independently created, they are not expected to be transient. For example, a ﬁle server could also provide WayStation services to foreign clients. However, an enduser machine is not an appropriate WayStation because
it may be turned off or disconnected at any time.
The decision of whether or not to use a WayStation is
client-driven. Clients estimate the quality of the network
between themselves and their current replica site . If
the network quality becomes poor, the client looks for
a WayStation close enough to improve matters, and initiates a replica on it. Such searches are carried out in
a neighborhood near the client through a process called
distance-based discovery . The WayStation informs
the server of initiation. This notiﬁcation is the point at
which the invalidation semantics on a client change. It is
an expensive operation, but it enables the server to drop
any outstanding callbacks for this client. Requiring callback breaks for wide-area clients would substantially penalize local-area clients.
The WayStation fetches objects on demand. We considered prefetching, but rejected it. It is not clear that inferred prefetching  can provide requests in time to be
useful given wide-area delays. We are also unwilling to
require that applications disclose prefetch requests .
Foreign clients must trust WayStations to hold cached
ﬁles securely and reconcile updates promptly. While one
cannot prevent WayStation misbehavior completely, we
plan to provide mechanisms to prevent exposure and reliably detect data modiﬁcation or repudiation of accepted
updates. This allows WayStation/client relationships to
be governed by a contract. Parties to a contract are not
prevented from breaching it; rather, the contract speciﬁes
penalties and remedies in the event of a breach. The details of our security architecture are beyond the scope of
this paper . We comment further on it in the context
of future work in Section 6. 3.3 Reconciliation: Managing Replicas Once a client has selected a WayStation, that WayStation
is treated as the client’s ﬁle server. Clients cache ﬁles
from WayStations, which manage callbacks for those
clients. Dirty ﬁles are written back to WayStations
on close. WayStations maintain update logs using a
mechanism similar to Coda’s client modify log . It
contains ﬁle stores, directory operations, and meta-data
updates. Redundant log records are removed via cancellation optimizations.
The server maintains an update log whenever one or
more WayStations hold replicas of that server. The server
also tracks which ﬁles are cached at the WayStation,
called the interest set. The bookkeeping for interest sets
is similar to that for callbacks, but the information is used
only during reconciliation.
Periodically, each WayStation reconciles its update log
with that of the server, exposing the updates made at each replica to the other one. To initiate a reconciliation,
the WayStation sends the server its update log plus a list
of ﬁles the WayStation has evicted from its cache. The
server removes the latter from the WayStation’s interest
set, and checks each WayStation log record to see if it is
serializable. If it is, the server invalidates the modiﬁed
object, and records the WayStation as the replica holding
The server responds with a set of ﬁles to invalidate, and
a set of ﬁles that are now in conﬂict, if any. The invalidation set consists of any updates accepted by the server—
whether from a client or through another WayStation—
that also reside in the WayStation’s interest set. Since the
WayStation will invalidate them, the server can remove
them from the interest set as well.
WayStations truncate their update logs on successful
reconciliation. However, since the server maintains a
merged log of all replica sites, it cannot truncate its log
immediately. The server maintains the last reconciliation
time for each WayStation holding a replica; let ØÓÐ ×Ø
be the earliest such time. The server need only hold log
records from ØÓÐ ×Ø forward, since all older records have
already been seen by all WayStations. Slow WayStations can force the server to keep an arbitrarily large log.
This policy is in contrast to that taken by Coda’s implementation of optimistic server replication. In Coda,
only the tails of large update logs are retained; objects
with discarded log records are marked in conﬂict and
will need manual repair. Given Coda’s presumption of
tightly-coupled and jointly-administered replicas, this is
an appropriate design choice. However, since WayStations are administered by independent entities, it would
be unwise to allow a WayStation’s absence to necessitate
many manual repairs.
If each WayStation reconciles at a constant rate, all updates are globally visible within twice the longest reconciliation period. The ﬁrst reconciliation invalidates the
server copy, and all other WayStation copies are invalidated during the next round. In order to provide clients
with well-bounded visibility, reconciliation must be a
lightweight operation. This is why reconciliations exchange only notiﬁcation of updates, not their contents.
Because sharing is rare, aggressively exchanging ﬁle
contents increases reconciliation time without improving client access time. Leaving shared accesses to pay
the full cost of wide-area network delays preserves performance, safety, and visibility for common-case operations. 3.4 Escrow: Providing Copy Semantics When a replica site needs the contents of an invalidated
ﬁle, what version does it obtain? The simplest approach
would be to provide the most recent version at the time WS 1 Server
fetch A WS 2
fetch A, B update A->A'
reconcile fetch B' read B'
read A reconcile This ﬁgure shows how fetching the most recent version
of a missing ﬁle can result in the propagation of out-oforder updates. WayStation 2 updates A then B. WayStation 1 sees the update to B ﬁrst, because A had already
been cached without an intervening reconciliation. Figure 1: Propagating Updates Out-Of-Order
it is requested; we call this last-version semantics. Unfortunately, last-version semantics allow WayStations to
see updates out of order. Instead, Fluid Replication’s
copy semantics guarantees that updates are seen in the
order they are made. We believe this is important, as
other studies have shown that users do not always understand the implications of complicated consistency mechanisms .
Figure 1 illustrates how out-of-order updates can arise
under last-version semantics. There are two ﬁles, and
, shared by clients at two WayStations, WS½ and WS¾ .
Assume that WS½ caches , and WS¾ caches both and
. Suppose that a client at WS¾ updates then , and
then reconciles. The server now knows about both new
versions held at WS¾ . If a client at WS½ were to referwithout an intervening reconence both ﬁles and
ciliation, it would see the old version of but the new
version of . When WS½ eventually reconciles, the update of would be known, but it would have been seen
out of order.
One solution would be to require clients to fetch all updated objects of interest whenever fetching any. Fluid
Replication eschews this approach, as it would result in
more work across the wide-area path between WayStation and server. Instead, we appeal to copy semantics.
Under copy semantics, WS½ would have read rather
than ¼ , because was the current version when WS½
last reconciled. This is despite the fact that was not in
WS½ ’s cache and the update, ¼ , had already been reconciled by WS¾ . Copy semantics guarantees that each WayStation see a complete preﬁx of all activity at each
WayStation, and therefore provides a consistent view of
the system. It is important to note that providing copy
semantics does not group related updates into atomic actions. In other words, copy semantics cannot prevent
seeing one update but not a related one. However, all
such mis-matches will be explicable in reference to “real
time”. In Bayou’s terminology, this provides monotonic
To provide copy semantics to a WayStation requesting a
ﬁle, Fluid Replication must supply the version known to
the server at the WayStation’s last reconciliation. This
version may reside on the server directly, or it may reside on the WayStation that had accepted it. Therefore,
a replica site that makes an update visible to another via
reconciliation must retain a copy of the update for as long
as the other replica might refer to it. We say that such
copies are held in escrow.
If a client at the server refers to a version escrowed at
some WayStation, the server back-fetches it; responsibility for the escrowed version passes from WayStation to
server. If a client at another WayStation references it, the
fetch request is ﬁrst sent to the server. The server obtains it from the escrowing WayStation—back-fetching
as before—and then forwards it on to the requesting
WayStation. Responsibility for escrow always migrates
from a WayStation to the server, never the reverse.
For escrow to be practical, we must prevent unbounded
storage growth. To see how this might occur, consider a
WayStation at which a ﬁle is updated between each reconciliation. Without knowing which versions are visible
to other replica sites, the WayStation would be forced to
hold all of them. The key to managing this growth is to
observe that only the most recent version visible to each
replica need be held in escrow. Any other versions are
said to be irrelevant, and can be safely pruned.
A simple way to prune versions is to track which are old
enough to be globally irrelevant. Recall that ØÓÐ ×Ø is
the latest time after which all WayStations have reconciled successfully. Sending this time to a WayStation
during a reconciliation allows the WayStation to prune
old versions. Let ½
Ò be the sequence of updates to
some ﬁle at a WayStation that were reconciled with the
be the most recent of those updates performed before ØÓÐ ×Ø . Clearly, all versions ½ ½
are irrelevant; each WayStation would request either
or some later version.
If all WayStations reconcile at the same rate, each
WayStation must keep at most one additional version of
a ﬁle in escrow. This is because each new reconciliation would ﬁnd that ØÓÐ ×Ø had advanced past the time
of that WayStation’s last reconciliation. A version enters escrow only when a more recent version is created; after reconciliation, ØÓÐ ×Ø advances past the new version,
rendering the old one irrelevant. Note that this scheme
allows some versions of a ﬁle stored at WayStations to
be discarded without ever being sent to the server. However, the ØÓÐ ×Ø mechanism guarantees that no replicas
still hold an active reference to that version.
While this pruning is helpful, it does not prevent the possibility of unbounded escrow space. Consider what happens if one WayStation is slow to reconcile. This single
WayStation prevents the advancement of ØÓÐ ×Ø , requiring the retention of all future versions at each WayStation. In effect, a single WayStation can hold all others
hostage. So, instead of sending only ØÓÐ ×Ø , the server
sends a sorted list of timestamps ØÓÐ ×Ø
ØÒ Û ×Ø ,
where each entry lists the time of last reconciliation by
some WayStation. If two or more visible versions lie between adjacent timestamps, only the last one is needed.
Under this scheme, a slow-to-reconcile WayStation requires only one additional copy in escrow.
The escrow mechanism depends on the fact that WayStations are closer in spirit to servers than they are to clients.
Because they hold versions in escrow for other replicas,
they must be reasonably available. An alternative design for Fluid Replication allows a client to act as its
own WayStation. While this approach would be simpler,
it can leave escrowed copies unavailable to other sites.
End-user machines are often unavailable for days at a
time, and may even disappear completely .
Under escrow, ﬁle updates can reside on WayStations
indeﬁnitely. One could argue that this is proper; if an
update is never needed elsewhere, it should never be
shipped from the WayStation. However, there are two
reasons to ship versions in advance of need. First, administrative tasks such as backup are greatly simpliﬁed
if ﬁles eventually migrate to the server. Second, migration limits the risk of unavailable escrowed copies and
data loss. Therefore, any version that has not been updated for one hour is sent to the server. We chose one
hour to take advantage of most cancellation opportunities , while balancing propagation overhead with the
need for availability.
It is important to note that copy semantics cannot guarantee in-order updates if those updates are destined for
more than one server. This is because there is no common synchronization point. A WayStation could provide
this guarantee through synchronous, atomic reconciliations across all of its servers. However, such steps introduce wide-area latencies into the critical path of operations, counter to Fluid Replication’s philosophy. Alternatively, the servers could synchronize amongst themselves, but with similar drawbacks. We expect that most
users will have one ”home” ﬁle system, obtaining the WayStation Server ate s com n. ns,
con Assume Success
Begin Escrow ns,
inv re xa
p Assume Success
Begin Escrow upd mit Server
co con Server Locked,
Cannot Abort WayStation Locked,
Can Abort WayStation xac tion (a) bilateral reconcliation (b) unilateral reconcliation This ﬁgure illustrates the difference between bilateral and unilateral reconciliations. Bilateral reconciliations require
three messages. Each replica must be locked during an expensive round trip, and WayStation failures can hold the
server hostage. In contrast, unilateral reconciliations optimistically assume that exposures are successful, and begin
escrow immediately. They require two messages rather than three, employ only short-term, local locks, and hold no
replicas hostage to failure. Figure 2: The Advantage of Unilateral Reconciliations beneﬁt of in-order updates while paying only the small
cost of local escrow. 3.5 Fault Tolerance in Reconciliation The simplest model for reconciliation is bilateral: the
atomic exchange of all update log records between a
WayStation and a server. Unfortunately, this simple
model is problematic in the face of node or network failures. Atomic exchanges require a two-phase commit
protocol. One node must prepare the transaction, agreeing to either commit or abort until the other party conﬁrms the result. In the meantime, the prepared node must
block many operations until the transaction completes.
The difﬁculty is caused by store operations during a bilateral reconciliation. These stores cannot be serialized
before the reconciliation. Doing so would require that
they had been in the reconciled update log, which is impossible. The stores cannot be serialized after the reconciliation either, since they may refer to a ﬁle that the
reconciliation will invalidate. Therefore, store operations
issued during a bilateral reconciliation must block until
it completes. In the presence of failures, stores may be
blocked indeﬁnitely. Put another way, bilateral reconciliation imposes wide-area networking costs on clients
even in the absence of sharing; this runs counter to Fluid
In light of these problems, we split a single, bilateral
reconciliation into two unilateral ones. These alternatives are illustrated in Figure 2. WayStations initiate reconciliations, as before. However, as soon as the rec- onciliation message is sent, the WayStation assumes it
will be received successfully. It can continue processing
client requests immediately, placing versions in escrow
as needed. The server likewise assumes success and begins escrow after sending its message. As a side effect,
the server’s message conﬁrms the exposures assumed by
the WayStation. The next WayStation request conﬁrms
the completion of the prior reconciliation.
There are several beneﬁts to this approach. WayStations block updates only while computing a reconciliation message; no locks are held across expensive
round trips. This appeals to the common case by placing potentially-visible versions in escrow immediately,
rather than waiting for conﬁrmation. Since the escrow
mechanism is needed to provide copy semantics, no additional complexity is required. Finally, the third reconciliation message—required to keep bilateral locking
times short—is implied by future messages. This piggybacking is similar to the process Coda uses to manage
updates to well-connected, replicated servers . Unlike Coda’s COP2 messages, conﬁrmation of reconciliations can be deferred indeﬁnitely. The only penalty of an
undetected failure is a larger escrow.
Unilateral reconciliations provide a good approximation
to the desired effect of bilateral ones. The server sees a
single, atomic event, while WayStations are able to allow clients to continue during the wide-area operation.
In addition to a potential increase in escrow size, there
is a slight widening of the conﬂict window, because concurrent writes are now allowed. Suppose that a WaySta- tion initiates a reconciliation, then accepts an update to
ﬁle that is later invalidated by the server. This causes
to be in conﬂict. With bilateral reconciliation, the update to would have been delayed until after invalidation
and then rejected, avoiding a conﬂict. However, given
the low incidence of write-shared ﬁles—particularly over
such short time frames—it is unlikely that such spurious
conﬂicts will occur in practice. stack, and returned objects must be copied off of the Java
stack. We were able to avoid making these copies by using preserialized objects, provided by the Jaguar package . Jaguar allows objects to be created outside the
Java VM, and still be visible from within. We used this
to pass objects, copy-free, between our C and Java code. 4 Implementation In evaluating Fluid Replication, we set out to answer the
following questions: Our Fluid Replication prototype consists of three components: a client cache manager, a server, and a WayStation. Each of these components is written primarily in
Java. There are several reasons for this, foremost among
them Java’s clean combination of thread and remote procedure call abstractions and the beneﬁts of writing code
in a type-safe language.
The bulk of the Fluid Replication client is implemented
as a user-level cache manager, supported by a small,
in-kernel component called the MiniCache . The
MiniCache implements the vnode interface  for Fluid
Replication. It services the most common operations for
performance, and forwards ﬁle operations that it cannot
satisfy to the user-level cache manager.
Calls are forwarded from the kernel to the cache manager across the Java Native Interface, JNI. The calls are
then satisﬁed by one of a set of worker threads, either
from the local disk cache or via Remote Method Invocation, RMI, to the appropriate replica. Fluid Replication
uses write-back caching; a close on a dirty ﬁle completes
in parallel with the store to the replica. The client supports dynamic rebinding of servers and WayStations for
migration; currently, our prototype migrates only on user
WayStations and servers share much of the same code
base, since their functionality overlaps. Data updates are
written directly to the replica’s local ﬁle system. Metadata is stored in memory, but is kept persistently. Updates and reconciliations are transactional, but we have
not yet implemented the crash recovery code. We use
Ivory  to provide transactional persistence in the Java
heap; used in this way, it is similar to RVM .
The decision to write the client in Java cost us some performance, and we took several steps to regain ground.
Our ﬁrst major optimization was to hand-serialize RMI
messages and Ivory commit records. The default RMI
skeleton and stub generator produced inefﬁcient serialization code, which we replaced with our own. This reduced the cost of a typical RMI message by 30%.
The second major optimization concerned the crossing
of the C-Java boundary. Each method call across this
boundary copies the method arguments onto the Java 5 Evaluation ¯
¯ Can Fluid Replication insulate clients from widearea networking costs?
What is the impact of sharing on performance?
How expensive is reconciliation?
Can Fluid Replication provide the consistency expected by local-area clients to those in the wide
How does escrow affect availability?
What are the storage costs of escrow? These questions concern the performance seen by individual clients and the behavior of the system as a whole.
To measure client performance, we subjected our prototype to a set of controlled benchmarks. We explored
system behavior through the use of trace-based simulation.
Our benchmarks ran on the testbed depicted in Figure 3. The WayStations are connected to the server via
a trace modulated network. Trace modulation performs
application-transparent emulation of a slower target network over a LAN . We have created modulation
traces that emulate the performance of a variety of different wide-area networking scenarios, listed in Table 1.
Latency numbers in these traces are in addition to baseline latency, while bandwidth speciﬁes the bottleneck capacity. All testbed machines run the Linux 2.2.10 kernel.
The server and WayStations have 550 MHz Pentium III
Xeon Processors, 256 MB of RAM, and 10K RPM SCSI
Ultra Wide disks. The clients are IBM ThinkPad 570s;
these machines have 366 MHz mobile Pentium IIs with
128 MB of memory.
For the trace-based studies, we collected traces comprising all activity on a production NFS server over
one week. This server holds 188 users’ home directories, plus various collections of shared data, occupying
48 GB. The users are graduate students, faculty, and staff
spread throughout our department. They come from a
variety of research and instructional groups, and have diverse storage needs. Generally speaking, the clients are
not mobile, so they may not be wholly representative of
our target domain. However, prior studies suggest that,
at the operation level captured by our traces, mobile and
desktop behavior are remarkably similar . Client WayStation Single Client Server with
Trace Modulation WayStation Client Two Clients (Sharing) This ﬁgure illustrates our benchmarking topology. Each
client is well connected to its WayStation, but trafﬁc between a WayStation and the server is subject to trace
modulation. Figure 3: Benchmark Topology
high latency Latency (ms)
10 Mb/s This table lists the parameters used in each of our trace
modulation scenarios. The local area scenario is the
baseline against which we compare. The next four were
obtained by measuring small ping and large ftp performance to four different sites. Bandwidth numbers are
increased over ftp throughput by 20% to account for
the difference in metric. The last two are used only to
determine trends as latency or bandwidth worsen, and
modify parameters orthogonally. Latencies are one-way;
bandwidths are symmetric. Table 1: Trace Modulation Parameters
Traces were collected using tcpdump on the network
segment to which the server was attached. These packet
observations were then fed into the nfstrace tool ,
which distilled the traces into individual fetch and store
operations. Note that this tool does not record operations
satisﬁed by a client’s cache. However, since NFS clients
do not use disk caches, this will overstate the amount
of read trafﬁc a Fluid Replication client would generate.
For the purposes of our analyses, we assume that each
client host resides on a separate WayStation. The traces
name 84 different machines, executing 7,980 read operations and 16,977 write operations. There are relatively
few operations because most of our client population did
not materially contribute to the total. Seven hosts account for 90% of all requests, and 19 hosts account for
99% of all requests. 5.1 Wide-Area Client Performance How effectively does Fluid Replication isolate clients
from wide-area networking costs? To answer this ques- tion, we compare the performance of Coda, AFS, and
Fluid Replication in a variety of networking conditions.
For Coda, we ran Coda 5.3.13 at the server and the client.
For AFS, we used OpenAFS, a descendant of AFS 3.6,
as the server, and Arla 0.35.3 as the client. To provide a
fair comparison, we set up our Coda volume on a single
server, rather than a replicated set.
Our benchmark is identical to the Andrew Benchmark  in form; the only difference is that we use the
gnuchess source tree rather than the original source tree.
Gnuchess is 483 KB in size; when compiled, the total
tree occupies 866 KB. We pre-conﬁgure the source tree
for the benchmark, since the conﬁguration step does not
involve appreciable trafﬁc in the test ﬁle system. Since
the Andrew Benchmark is not I/O-bound, it will tend
to understate the difference between alternatives. In the
face of this understatement, Fluid Replication still outperforms the alternatives substantially across wide-area
We tested each ﬁle system with both cold and warm
caches. In the case of AFS and Coda, a “warm cache”
means that the clients already hold valid copies of the
gnuchess source tree. In the case of Fluid Replication,
the source tree is cached on the WayStation.
Figures 4 compares the total running times of the Fluid
Replication, Coda, and AFS clients under different network environments. Figure 4(a) gives the performance
with a cold cache, and Figure 4(b) shows it with a warm
cache. Each experiment comprises ﬁve trials, and the
standard deviations are less than 2% of the mean in all
With a cold cache and a well-connected server, Coda
and AFS outperform Fluid Replication. We believe this
is due to our choice of Java as an implementation language and the general maturity level of our code. We see
no fundamental reason why Fluid Replication’s performance could not equal Coda’s. They have identical client
architectures, and the client/WayStation interactions in
Fluid Replication are similar to those between client and
server in Coda.
As the network conditions degrade, the cold-cache times
of Coda and AFS rapidly increase while those of Fluid
Replication increase slowly. All systems must fetch
source tree objects from the server, and should pay similar costs to do so. The divergence is due to the systems’
different handling of updates. In Fluid Replication, all
updates go only to the nearby WayStation. In Coda and
AFS, however, updates must go across the wide area to
With a warm cache, the total running time of Fluid Replication remains nearly constant across all network environments. This is because the updates are propagated 120 Total Time (s) 100 Trace
High Latency fr-cold
local small medium large intercontinental Network (a) cold cache
120 Total Time (s) 100 fr-warm
1.007 This ﬁgure shows the normalized running time for Fluid
Replication over each of the networking scenarios in Table 1. The local-area case served as the performance
baseline in each case. When the WayStation cache is
cold, performance decreases as wide-area networking
costs increase. However, no consistent trend exists when
the WayStation cache is warm. Table 2: Fluid Replication over Wide-Area Networks 80 local Cold
1.241 medium large intercontinental Network (b) warm cache
This ﬁgure compares the total running times of Coda,
AFS and Fluid Replication under a variety of network
environments. The upper ﬁgure gives the results for a
cold cache, and the lower ﬁgure shows them for a warm
cache. With a cold cache, Fluid Replication is least affected by network costs among three systems. With a
warm cache, the performance of Fluid Replication does
not appreciably degrade as network costs increase. Figure 4: Client Performance
only to the WayStation; invalidations are sent to the
server asynchronously. The running time of AFS and
Coda increases as the network degrades. They must
propagate updates to the server during the copy phase
and write object ﬁles back to the server during the make
phase. The Coda client never entered weakly-connected
mode, since the bandwidths in our sample traces were
well above its threshold of 50 KB/s. Had Coda entered
weakly-connected mode, its performance would be identical to Fluid Replication’s, but its updates would be neither safe nor visible for many minutes.
Table 2 shows the normalized running time of Fluid
Replication, with the local-area case serving as the baseline. While the running time increases by as much as
24.1% when the cache is cold, it remains nearly constant
when the cache is warm. Looking only at the four scenarios generated by ping and ftp experiments, there appears to be a slight growth trend, but it is within observed
variance. Only the results for Small and Intercontinental
are not identical under the Ø-test; all other pairs are sta- tistically indistinguishable. To conclusively rule out a
true trend, we also ran the warm-cache Fluid Replication
experiment across two more-demanding networking scenarios. The ﬁrst decreased bandwidths to 56 Kb/s, but
added no additional latency. The second increased oneway latency to 200 ms, but placed no additional bandwidth constraints. In both cases, Fluid Replication’s running time was less than that for the Intercontinental trace.
Therefore, we conclude that that Fluid Replication’s update performance does not depend on wide-area connectivity. Of course, this is due to deferring the work of update propagation and the ﬁnal reconciliation — neither of
which contribute to client-perceived performance. Sections 5.4 and 5.5 quantify any reduction in consistency
or availability due to deferring work. 5.2 Costs of Sharing Our next task is to assess the potential impact that deferred propagation has on sharing between wide-area
clients. We devised a benchmark involving two clients
sharing source code through a CVS repository. In the
benchmark, clients C1 and C2 are attached to WayStations W1 and W2, respectively. In the ﬁrst phase, C1
and C2 each check out a copy of the source tree from
the repository. After changing several ﬁles, C1 commits
those changes to the repository. Finally, C2 updates its
source tree. Both the repository and working copies reside in the distributed ﬁle system. When the benchmark
begins, both clients have the repository cached.
We used the Fluid Replication source tree for this benchmark. The changes made by C1 are the updates made
to our source tree during the four of the busiest days
recorded in our CVS log. At the beginning of the trial,
the source tree consists of 333 ﬁles totaling 1.91 MB.
The committed changes add four new ﬁles, totaling seven
KB, and modify 13 others, which total 246 KB. We reran this activity over the two-client topology of Figure 3. 50 Time (s) 40 every update to the server. Furthermore, these updates
break callbacks on the far client. fr-commit
local small medium large intercontinental Network (a) commit
50 Time (s) 40 fr-update
local small medium large intercontinental Network (b) update
Our sharing benchmark replayed sharing from the CVS
log of our Fluid Replication source tree. The benchmark
consisted of ﬁve phases: a checkout at C1, a checkout at
C2, an edit at C1, a commit at C1, and an update at C2.
This ﬁgure shows the time to complete the commit and
update phases for AFS, Coda, and Fluid Replication. Figure 5: Sharing Benchmark
Figure 5 shows our results for the commit and update
phases, run over Fluid Replication, AFS, and Coda; each
bar is the average of ﬁve trials. We do not report the
results of the checkout and edit phases, because they exhibit very little sharing.
Unsurprisingly, the cost of committing to the repository
was greater for AFS and Coda than for Fluid Replication.
This is because Fluid Replication clients ship updates to
a nearby WayStation. AFS and Coda must ship data and
break callbacks over the the wide area.
What is surprising is that the update phase is also more
costly for AFS and Coda. One would think that Fluid
Replication would perform poorly, since ﬁle data has to
traverse two wide area paths: from the ﬁrst WayStation
to the server, and then to the second WayStation. The unexpected cost incurred by AFS and Coda stems from the
creation of temporary ﬁles. The latter are used to lock the
repository and back up repository ﬁles in case of failure.
WayStations are able to absorb these updates, and later
optimize them away, since they are all subject to cancellation. AFS and Coda clients, on the other hand, send It is important to note that hiding the creation of temporary locking ﬁles may improve performance, but it renders the intended safety guarantees useless. The best that
Fluid Replication can offer in the face of concurrent updates is to mark the shared ﬁle in conﬂict; the users must
resolve it by hand. Coda would be faced with the same
dilemma if it had entered weakly-connected mode during this benchmark. However, we believe that in practice, optimism is warranted. Even in the case of CVS
repositories, true concurrent updates are rare. Our own
logs show that commits by different users within one
minute occurred only once in over 2,100 commits. A
commit followed by an update by different users within
one minute happened twice.
This benchmark illustrates the cost of obtaining deferred
updates. However, in practice, these ﬁles are likely to
have migrated to the server. Our logs show that the median time between commits by different users was 2.9
hours, and that the median time between a commit and
an update by different users was 1.9 hours. This would
provide ample opportunity for a WayStation to asynchronously propagate shared data back to the server before it is needed. 5.3 Reconciliation Costs To be successful, Fluid Replication must impose only
modest reconciliation costs. If reconciliations are expensive, WayStations would be able to reconcile only infrequently, and servers could support only a handful of
WayStations. To quantify these costs, we measured reconciliations, varying the number of log records from 100
to 500. To put these sizes in context, our modiﬁed Andrew Benchmark reconciled as many as 148 log records
in a single reconciliation.
Figure 6 shows the reconciliation time spent at the server
as the number of log records varies; each point is the
average of ﬁve trials. This time determines the number of WayStations a server can handle in the worst case.
Server-side time increases to just under 1.1 seconds for
500 log records. In practice we expect the costs to be
much smaller. The week-long NFS trace would never
have generated a reconciliation with more than 64 log
Figure 7 shows the total time for a reconciliation measured at a WayStation as the number of log records
varies. This includes the server-side time, RMI overheads, and network costs. The total reconciliation times
for the local area, small distance, and medium distance
traces do not vary signiﬁcantly. This means that, at these
speeds, bandwidth is not the limiting factor. Proﬁling of
the reconciliation process suggests that RMI—even with 0.25 1.2 % of All Operations Time (s) 1.0
0 0 100 200 300 400 500 0.25 1 4 Number of Log Records This ﬁgure shows the reconciliation times spent at a
server as the number of log records varies. These times
determine the number of WayStations a server can handle. Figure 6: Reconciliation Time at Sever
local 3.0 Time (s) 2.5
100 200 300 400 500 Number of Log Records This ﬁgure shows the total reconciliation times in seconds as the number of log records varies. These times
determine the upper limit on how frequently a WayStation can reconcile with the server. For all sizes, they are
much shorter than our default reconciliation period of 15
seconds. Figure 7: Reconciliation Time at WayStation
our hand-optimized signatures—is the rate-limiting step.
However, at 1 Mb/s, the bandwidth of the large distance
and intercontinental traces, the bottleneck shifts to networking costs. In any event, the reconciliation times are
much shorter than our default reconciliation period of 15
seconds, allowing the WayStation to reconcile more frequently if sharing patterns warrant. 5.4 64 256 1024 4096 This ﬁgure shows the percentage of operations that
caused sharing. The x-axis shows the time between uses
of a shared objects in minutes, and the y-axis shows the
percentage of total operations that exhibited sharing. The
top part of bar shows the read-after-write sharing; the
bottom part shows the write-after-write. Only 0.01% of
all operations caused sharing within 15 seconds. Just
over 0.6% of all operations exhibited any form of sharing. Figure 8: Sharing 1.5 0 16 Time between Uses (minutes) Consistency The consistency offered by Fluid Replication depends on
two factors: the frequency of reconciliation and the time
between uses of a shared object. Section 5.3 quantiﬁed
the former. In this section, we address the incidence of
sharing observed in a real workload.
To determine how often sharing happens and the time
between shared references, we examined our week-long
NFS client traces. For this analysis, we removed references to user mail spools from our traces. A popular mail client in our environment uses NFS rather than IMAP
for mail manipulation. Many users run these clients on
more than one machine, despite NFS’s lack of consistency guarantees , generating spurious shared references. Since we would expect mobile users to use IMAP
instead, we excluded these references.
Figure 8 shows the percentage of all references to objects written previously at another replica site. The top
part of each bar shows read-after-write sharing, and the
bottom part shows write-after-write. As expected, sharing is not common, especially over short periods of time.
Only 0.01% of all operations caused sharing within 15
seconds. The total fraction of references that exhibited
sharing during the week was just over 0.6% of all operations. Note that these numbers are pessimistic, as we
have assumed that each client uses a distinct WayStation.
The graph shows some interesting periodic behavior; unfortunately, with network-level traces, we are unable to
identify the processes causing it. 5.5 Availability Because WayStations do not propagate ﬁle contents to
the server immediately, a failed WayStation could keep
a client from retrieving a needed update. To gauge
how likely such a scenario might be, we fed our NFS
traces into a Fluid Replication simulator. We augmented the trace with WayStation reconciliations and
failure/recovery pairs. Reconciliations were scheduled
every 15 seconds from the WayStation’s ﬁrst appearance
in the trace. We assume a mean time to failure of 30 days
and a mean time to repair of one hour, both exponentially distributed. These parameters were chosen to rep- Number of Reconciliations 10
8 MBytes 7
3 4500 total reconciliations with non-empty escrows: 24843 4000
1K 2 4K 64K 256K 1M 4M 16M Escrow Size (byte) 1
1 2 3
Days 4 5 6 This ﬁgure shows the size of global escrow. The x-axis
shows time in days from the beginning of the NFS trace,
and the y-axis shows size of escrow. Figure 9: Global Escrow Size
resent typical uptimes for carefully administered server
machines; they are the same as those used in the xFS
study . Note that we did not model server failures; our
intent is to explore Fluid Replication’s additional contribution to failures.
For each trial of the simulation, we varied the random
seed controlling failures and repairs. We then counted
the number of back-fetches that fail due to the unavailability of a WayStation. We took over ﬁve million trials
of the experiment in order to provide reasonable conﬁdence intervals for a very small mean.
Recall that during our trace there were 24,957 total requests made of the WayStations. Of those, 243 requests
required a back-fetch from one WayStation to another.
We ﬁrst simulated Fluid Replication when updates that
had not been overwritten or back-fetched were propagated by the WayStation after 1 hour. Across ﬁve million trials, 187,376 requests failed, for an average of
¢ ½¼ failures per operation, with a 90% conﬁdence interval of ¦ ¿ ¢ ½¼ ; this is equivalent to ﬁve
nines’ availability. Expressed in time, we observed an
average of 0.037 failures per week, or roughly one failed
access every 27 weeks for our client population. 5.6 16K Costs of Escrow Our ﬁnal question concerns the storage costs of escrow.
To determine them, we fed the week-long NFS traces
into our simulator. Each WayStation in the trace reconciles every 15 seconds from the moment it is ﬁrst named,
as before. After an update has been reconciled, it is
marked eligible for escrow. Any subsequent write to the
ﬁle checks for this mark, and, in its presence, preserves
the old version in case it is needed. Updates are removed
from escrow as described in Section 3.4. This ﬁgure shows the distribution of non-empty escrow
sizes. The x-axis shows time in days from the beginning
of the NFS trace, and the y-axis shows the number of
reconciliations resulting in an escrow of that size or less. Figure 10: Non-Empty Escrow Distribution
Figure 9 shows how global escrow size changes with
time. The x-axis plots days from the beginning of the
trace and the y-axis shows the size of escrow in MB.
Global escrow size was almost always zero, and never
exceeded 10 MB over the course of the week. The large
variance matches the burstiness expected of update activity. Over the lifetime of our trace, WayStations collectively handle over 280 MB; compared to this, escrow
requirements are modest.
Figure 10 gives a more detailed picture of escrow sizes
over time. This histogram plots reconciliations resulting in non-empty escrows, showing the frequency with
which each size was seen. Almost 89% of all non-empty
escrows are 2MB or smaller. 6 Future Work
There are three main tasks ahead of us. First, we plan
to use network estimation  and distance based discovery  to automate the WayStation migration and
location processes. We have developed a technique that
estimates network performance along the path between a
client and a server using observations of request/response
trafﬁc. This estimator, borrowing techniques from statistical process control, follows underlying shifts in network performance while ﬁltering out observed noise.
In effect, the estimator adapts its behavior to prevailing networking conditions, selecting for agility or stability as appropriate. This estimator, combined with a
cost-beneﬁt analysis of WayStation migration, guides the
search for a new replica site when the client moves too
far from the old one.
Our second task is to address the need for trust between clients and WayStations . Before a client will
agree to use a WayStation, it must be assured of the privacy and integrity of cached data, and of non-repudiation
of updates. Preventing exposure through encryption is straightforward, though managing keys can be subtle. It
is also easy to reliably detect unauthorized modiﬁcations
using cryptographic hashes. However, guaranteeing that
a WayStation will correctly forward updates is difﬁcult,
if not impossible.
Rather than attempt to prove a WayStation trustworthy a
priori, we plan to provide a receipt mechanism that enables a client to prove that a WayStation did not properly
forward an update. When a client ships a version of a
ﬁle to the WayStation, it receives a cryptographicallysigned receipt for that update. The client can later check
whether that version was properly reconciled; lack of
reconciliation provides evidence of update repudiation.
Receipts can be optimized away much as updates can,
but clients must retain the authority to apply such optimizations. Furthermore, clients retain the latest version
of any updated ﬁle until that version is known to either
reside on the server or have been invalidated; the space
required to do so is modest . This guarantees that
the latest version of a ﬁle is known to reside on either a
client or the server, in addition to any untrusted WayStation copies. Thus, a WayStation that disappears from the
system cannot take with it the only current version of a
ﬁle, though earlier, escrowed versions may vanish.
Our third goal is to gain more experience with Fluid
Replication and its use. While the NFS traces are informative, they do not capture precisely how users will
use Fluid Replication. We plan to deploy a server in
our department for day-to-day storage requirements, and
provide users with WayStations and wireless gateways at
home. This will allow client laptops to seamlessly migrate between locales, giving us valuable insights into
the system’s use. 7 Conclusion
As mobile clients travel, their costs to reach back to home
ﬁling services change. To mask these performance problems, current ﬁle systems reduce either safety, visibility,
or both. This is a result of conﬂating safety and visibility into a single mechanism. They have different requirements, and so should be provided through different
Fluid Replication separates the concerns of safety and
visibility. While traveling, a mobile client associates itself with a nearby WayStation that provides short-term
replication services for the client’s home ﬁle system. Updates are sent to the nearby WayStation for safety, while
WayStations and servers frequently exchange knowledge
of updates through reconciliation to provide visibility.
Reconciliation is inexpensive and wait-free. WayStations retain copies of advertised updates in escrow until
they are irrelevant to other replica sites. An analysis of trafﬁc in a production NFS server validates our design decisions. A modest reconciliation interval of ﬁfteen seconds limits the rate of stale reads or
conﬂicting writes to 0.01%. Reserving 10 MB for escrow space—a small fraction of disk capacity—is sufﬁcient in the worst case. Measurements of a Fluid Replication prototype show that it isolates clients from most
wide-area networking costs. Update trafﬁc is not affected
at bandwidths as low as 56 Kb/s or latencies as high as
200 ms. These gains offset increased sharing costs, even
for workloads with substantial degrees of sharing. Escrow space requirements are modest: at its peak, it is
less than 5% of the total data cached at WayStations. Despite the fact that update propagation is deferred, availability does not suffer. Our traced clients could expect
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This note was uploaded on 11/12/2011 for the course CE 726 taught by Professor Staf during the Spring '11 term at SUNY Buffalo.
- Spring '11