Satori: Enlightened page sharing
Grzegorz Miłoś, Derek G. Murray, Steven Hand
University of Cambridge Computer Laboratory
Cambridge, United Kingdom
Michael A. Fetterman
Bedford, Massachusetts, USA
We introduce Satori
, an efficient and effective system for
sharing memory in virtualised systems. Satori uses
in guest operating systems to detect sharing
opportunities and manage the surplus memory that results from
sharing. Our approach has three key benefits over existing systems: it
is better able to detect short-lived sharing opportunities, it is
efficient and incurs negligible overhead, and it maintains performance
isolation between virtual machines.
We present Satori in terms of hypervisor-agnostic design decisions,
and also discuss our implementation for the Xen virtual machine
monitor. In our evaluation, we show that Satori quickly exploits up to
of the maximum possible sharing with insignificant
performance overhead. Furthermore, we demonstrate workloads where the
additional memory improves macrobenchmark performance by a factor of
An operating system can almost always put more memory to good use. By
adding more memory, an OS can accommodate the working set of more
processes in physical memory, and can also cache the contents of
recently-loaded files. In both cases, cutting down on physical I/O
improves overall performance. We have implemented Satori, a
novel system that exploits opportunities for saving memory when
running on a virtual machine monitor (VMM). In this paper, we explain
the policy and architectural decisions that make Satori efficient and
effective, and evaluate its performance.
Previous work has shown that it is possible to save memory in
virtualised systems by sharing pages that have
identical  and/or
similar  contents. These systems were designed for
unmodified operating systems, which impose restrictions on the sharing
that can be achieved. First, they detect sharing opportunities by
periodically scanning the memory of all guest VMs. The scanning rate
is a trade-off: scanning at a higher rate detects more sharing
opportunities, but uses more of the CPU. Secondly, since it
overcommits the physical memory available to guests, the VMM must be
able to page guest memory to and from disk, which can lead to
We introduce enlightened page sharing as a collection of techniques for
making informed decisions when sharing memory and distributing the benefits.
Several projects have shown that the performance of a guest OS running on a VMM
improves when the guest is modified to exploit the virtualised
environment [1,25]. In Satori, we add two main
enlightenments to guests. We modify the virtual disk subsystem, to
implement sharing-aware block devices: these detect sharing opportunities
in the page cache immediately as data is read into memory. We also add a
repayment FIFO, through which the guest provides pages that the VMM can
use when sharing is broken. Through our modifications, we detect the majority of
sharing opportunities much sooner than a memory scanner would, we obviate the
run-time overhead of scanning, and we avoid paging in the VMM.
We also introduce a novel approach for distributing the benefits of
page sharing. Each guest VM receives a sharing entitlement that
is proportional to the amount of memory that it shares with other
VMs. Therefore, the guests which share most memory receive the
greatest benefit, and so guests have an incentive to share. Moreover,
this maintains strong isolation between VMs: when a page is unshared,
only the VMs originally involved in sharing the page are affected.
When we developed Satori, we had two main goals:
The rest of this paper is organised as follows. We begin by discussing
the issues of memory management in both operating systems and virtualised
platforms (Section 2). We then survey related systems
(Section 3). We present Satori in two parts: first, we justify
the major design decisions that differentiate Satori from other systems
(Section 4), then we describe how we implemented a prototype of
Satori for the Xen VMM (Section 5). Finally, we perform a thorough
evaluation of Satori's performance, including its effectiveness at finding
sharing opportunities and its impact on overall performance
- Detect short-lived sharing:
- We show in the evaluation that the
majority of sharing opportunities are short-lived and do not persist
long enough for a memory scanner to detect them. Satori detects
sharing opportunities immediately when pages are loaded, and quickly
passes on the benefits to the guest VMs.
- Detect sharing cheaply:
- We also show that Satori's
impact on the performance of a macrobenchmark--even without the
benefits of sharing--is insignificant. Furthermore, when sharing
is exploited, we achieve improved performance for some
macrobenchmarks, because the guests can use the additional memory to
cache more data.
The problem of memory management has a long history in operating systems and
virtual machine monitors. In this section, we review common techniques for
managing memory as a shared resource (§ 2.1). We then
describe the relevant issues for page sharing in virtual machine monitors
(§ 2.2). Finally, we describe how paravirtualisation is
used to improve performance in virtualised systems (§ 2.3).
Physical memory is a scarce resource in an operating system. If more
memory is available, it can be put to good use, for example by
obviating the need to swap pages to disk, or by caching
recently-accessed data from secondary
storage. Since memory access is several orders of magnitude faster
than disk access, storing as much data as possible in memory has a
dramatic effect on system performance.
Memory resource management was first formalised for operating systems
by Denning in 1968, with the introduction of the working set
model . The working set of a process at time is
the set of pages that it has referenced in the interval
. This is a good predictor of what pages should be maintained in
memory. Pages can then be allocated to each process so that its
working set can fit in memory.
Since it is challenging to calculate the working set and
parameter exactly, an alternative approach is to monitor the page
fault frequency for each process . If a
process causes too many page faults, its allocation of pages is
increased; and vice versa. This ensures acceptable progress for all
OS-level approaches are inappropriate for a virtualised system. One of
the key benefits of virtualisation is that it provides resource isolation
between VMs. If the size of a VM's working set or its page fault rate is allowed
to determine its memory allocation, a malicious VM can receive more than its
fair share by artificially inflating either measure. Instead, in our approach,
we give a static allocation of physical memory to each VM, which provides strong
performance isolation . As we describe in
§ 4.2, our system provides surplus memory to VMs that
participate in sharing. Our approach follows previous work on
self-paging, which required each application to use its own resources
(disk, memory and CPU) to deal with its own memory faults .
1 Virtual memory management
A conventional operating system expects to own and manage a range of
contiguously-addressed physical memory. Page tables in these systems translate
virtual addresses into physical addresses. Since virtualisation
can multiplex several guest operating systems on a single host, not all guests
will receive such a range of physical memory. Furthermore, to ensure isolation,
the VMM's and guests' memory must be protected from each other, so the VMM must
ensure that all updates to the hardware page tables are valid.
Therefore, a virtualised system typically has three classes of
address. Virtual addresses are the same as in a conventional
OS. Each VM has a pseudo-physical address space, which is
contiguous and starts at address zero. Finally, machine
addresses refer to the physical location of memory in hardware. A
common arrangement is for guests to maintain page tables that
translate from virtual to pseudo-physical addresses, and the VMM to
maintain separate shadow page tables that translate directly
from virtual addresses to machine
addresses . A more recent approach is to use
additional hardware to perform the translation from pseudo-physical
addresses to machine addresses [10,19]. Finally, it
is also possible to modify the OS to use machine addresses and communicate
with the VMM to update the hardware page tables
Pseudo-physical addresses provide an additional layer of indirection that makes
it possible to share memory between virtual machines. Since, for each VM, there
is a pseudo-physical-to-machine (P2M) mapping, it is possible to make several
pseudo-physical frame numbers (PFNs) map onto a single machine
frame number (MFN). Therefore, if two VMs each have a page with the same
contents, the VMM can update the P2M mapping and the shadow page tables to make
those pages use the same machine frame. We discuss how other systems detect
duplicates in Section 3, and the Satori approach in
If two VMs share a page, an attempt to write to it must cause a page
fault. This is achieved by marking the page read-only in the shadow page table.
Such a page fault is called a copy-on-write fault. When this occurs, the
VMM handles the fault by allocating a new frame and making a private copy of the
page for the faulting guest. It also updates the P2M mapping and shadow page
tables to ensure that the guest now uses the private copy.
A consequence of page sharing is that the memory used by a VM can both
dynamically decrease (when a sharing opportunity is exploited) and
dynamically increase (when sharing is broken). This presents a
resource allocation problem for the VMM. A conventional operating
system does not have fine-grained, high-frequency mechanisms to deal
with memory being added or removed at run time (Memory hotplug
interfaces are unsuitable for frequent, page-granularity addition and
removal ). Therefore, one option is to use a
balloon driver in each guest, which pins physical memory within
a guest and donates it back to the
VMM [1,23]. The ``balloon''
can inflate and deflate, which respectively decreases and increases
the amount of physical memory available to a given VM.
However, a balloon driver requires cooperation from the guest: an alternative is
host paging, whereby the VMM performs page replacement on guests'
pseudo-physical memory [4,23]. Host paging
is expensive, because a VM must be paused while evicted pages are faulted in,
and even if the VMM-level and OS-level page replacement policies are perfectly
aligned, double paging (where an unused page must be paged in by the VMM when
the OS decides to page it out) negatively affects performance. We deliberately
avoid using host paging, and use a combination of the balloon driver (see
§ 4.2) and volatile pages (see
§ 4.3) to vary the memory in each guest dynamically.
Collaborative memory management (CMM) attempts to address the
issue of double paging . This system was
implemented for Linux running on IBM's z/VM hypervisor for the zSeries
architecture. In CMM, the guest VM provides hints to the VMM that suggest what
pages are being used, and what pages may be evicted with little penalty. In
Satori, we use part of this work for a different purpose: instead of using hints
to improve a host pager, we use them to specify pages which may be reclaimed
when sharing is broken (see § 5.3).
2 Memory virtualisation and sharing
Our approach to memory sharing is based on enlightenments, which
involve making modifications to operating systems in order to achieve the best
performance in a virtualised environment; in this paper we use the
terms ``enlightenment'' and ``paravirtualisation'' interchangeably.
Operating systems have been modified to run on VMMs for almost as long as VMMs
have existed: the seminal VM/370 operating system employs handshaking to
allow guests to communicate with the VMM for efficiency
reasons . ``Paravirtualisation'' was coined for the Denali
VMM , and Xen was the first VMM to run paravirtualised commodity
operating systems, such as Linux, BSD and Windows . Xen
showed that by paravirtualising the network and block devices, rather than
relying on emulated physical hardware, it was possible to achieve near-native
More extreme paravirtualisation has also been proposed. For example, Pfaff
et al. designed Ventana as a virtualisation-aware
file system, to replace virtual block devices as the storage
primitive for one or more VMs . This design concentrates
on adding functionality to the file system--for example versioning,
isolation and encapsulation--and considers sharing from the point of
view of files shared between users. It does not specifically address
resource management or aim to improve performance. Our
approach is orthogonal to Ventana, and similar memory sharing benefits
could be achieved with a virtualisation-aware file system. Indeed,
using a system like Ventana would probably make it easier to identify
candidates for sharing, and improve the overall efficiency of our
Other systems, such as VMware ESX
Server  and the Difference
Engine  have a design goal of supporting unmodified
guest OSs. In contrast, we have concentrated on paravirtualised guests
for two reasons. First, there is an increasing trend towards
enlightenments in both Linux and Microsoft Windows operating
systems [18,20]. Secondly, we
believe that where there is a compelling performance benefit in using
enlightenments, the necessary modifications will filter down into the
vanilla releases of these OSs.
Waldspurger described a broad range of memory management techniques
employed in the VMware ESX Server hypervisor, including page
sharing . In VMware ESX Server, page
sharing opportunities are discovered by periodically scanning the
physical memory of each guest VM, and recording fingerprints of each
page. When the scanner observes a repeated fingerprint, it compares
the contents of the relevant two pages, and shares them if they are
identical. In the same work, Waldspurger introduced the balloon driver
that is used to alter guest memory allocations. However, since VMware
ESX Server is designed to run unmodified guest operating systems, it
must also support host paging. In Satori, we avoid host paging because
of its negative performance impact (see § 2.2),
and avoid memory scanning because it does
not detect short-lived sharing opportunities (see
A contemporary research project has added page sharing to the Xen
Virtual Machine Monitor. Vrable et al. began this effort
with Potemkin , which uses flash cloning and
delta virtualization to enable a large number of
mostly-identical VMs on the same host. Flash cloning creates a new VM
by copying an existing reference VM image, while delta virtualization
provides copy-on-write sharing of memory between the original image
and the new VM. Kloster et al. later extended this work with
a memory scanner, similar to that found in VMware ESX
Server . Finally, Gupta et
al. implemented the Difference Engine, which uses patching and
compression to achieve greater memory savings than sharing alone. We
have implemented Satori on Xen in a parallel effort, but we use guest
OS enlightenments to reduce the cost of duplicate detection and memory
The Disco VMM includes some work on transparent page
sharing . In Disco, reading from a special
copy-on-write disk involves checking to see if the same block
is already present in main memory and, if so, creating a shared
mapping to the existing page. We apply a similar policy for duplicate
detection, as described in § 4.1. However, we
also implement content-based sharing for disk I/O
(§ 5.2), so it is not necessary to use
copy-on-write disks, and furthermore we can exploit identical blocks within
the same disk.
3 Related Work
In this section, we present the major design decisions that
differentiate Satori from previous work on page
sharing [23,4]. Figure 1
shows the life-cycle of a page that participates in sharing. This
diagram raises three key questions, which we address in this section:
4 Design decisions
We have taken care to ensure that our answers to the above questions are
hypervisor-agnostic and may be implemented together or individually. Although
our prototype uses the Xen VMM (see Section 5), these techniques should
also be useful for developers of other hypervisors. In particular, our duplicate
detection and savings distribution policies could be implemented without
modifying core OS components in the guest VMs. However, by enlightening the
guest OS, it is possible to achieve better performance, and we contend that our
techniques are best implemented as a whole.
- How are duplicates detected?
- We use sharing-aware
block devices as a low-overhead mechanism for detecting duplicate
pages. Since a large amount of sharing originates within the page
cache, we monitor data as it enters the cache (§ 4.1).
- How are memory savings distributed?
- When identical pages
are discovered, these can be represented by a single physical page,
and pages are saved. We distribute these savings
to guest VMs in proportion with their contribution towards
sharing (§ 4.2).
- What if sharing is broken?
- Shared pages are
necessarily read-only. When a guest VM attempts to write to a shared
page, the hypervisor makes a writable private copy of the
page for the guest. We require that the guest itself provides a list
of volatile pages that may be used to provide the necessary
memory for private copies. In addition, we obviate the need for
copying in certain cases (§ 4.3).
In order to exploit page sharing, it is necessary to detect duplicate
pages. As described in § 2.2, the most common
approach to this problem is to scan the memory of all guest VMs
periodically, and build up a list of page fingerprints that can be
used to compare page contents. In this subsection, we propose
sharing-aware block devices as a more efficient alternative. We
discuss the problems with the scanning-based approach, and explain why
the block interface is an appropriate point at which to detect
As we show in § 6.1, many sharing opportunities
are short-lived, and yet these provide a large amount of overall
sharing when taken as a whole. In principle, the memory
scanning algorithm is exhaustive and all duplicates will eventually be
found. However, in practise the rate of scanning has to be capped: in
the extreme case, each memory write would trigger fingerprint
re-computation. For example in VMware ESX Server the default memory
scan frequency is set to once an hour, with a maximum of six times
per hour . Therefore, the
theoretical mean duplicate discovery time for the default
setting is 40min, which means that short-lived sharing opportunities
will be missed. (We note that there are at
least three relevant configuration options: the scan period, scan
throughput (in MB per second per GHz of CPU), and maximum scan rate
(in pages per second). In our evaluation
(§ 6.1), the ``aggressive'' settings for
VMware use the maximum for all three of these parameters.)
When an operating system loads data from disk, it is stored in the
page cache, and other researchers have noted that between and
of shareable pages between VMs are part of the page
cache . For example, VMs based on the same
operating system will load identical program binaries, configuration
files and data files. (In these systems, the kernel text will also be
identical, but this is loaded by Xen domain builder (bootloader), and
does not appear in the page cache. Though we do not implement it here,
we could modify the Xen domain builder to provide sharing hints.)
The efficacy of sharing-aware block devices relies on the observation
that, for the purpose of detecting duplicates, a good image of the
page cache contents can be built up by observing the content of disk
reads. While this approach does not capture any subsequent in-memory
writes, we do not expect the sharing potential of dirty pages to be
high. Since a VMM uses virtual devices to represent block devices, we
have a convenient place to intercept block reads. We describe our
implementation of sharing-aware block devices in
The situation improves further if several guests share a common base
image for their block device. When deploying a virtualised system, it
is common to use a single substrate disk for several VMs, and
store VM-private modifications in a copy-on-write overlay file. If a
guest reads a block from the read-only substrate disk, the block
number is sufficient to identify it uniquely, and there is no need to
inspect its contents. This scheme has the additional advantage that
some reads can be satisfied without accessing the underlying physical
Previous work on page sharing emphasises zero pages as a large source
of page duplicates. Clearly, these pages would not be found by
block-device interposition. However, we take a critical view of zero-page
sharing. An abundance of zero pages is often indicative of low memory
utilisation, especially in operating systems which implement a
scrubber. We believe that free page sharing is usually counterproductive,
because it gives a false sense of memory availability. Consider the
example of a lightly loaded VM, in which of pages are zero pages.
If these pages are reclaimed and used to run another VM, the original
VM will effectively run with of its initial allocation. If this
is insufficient to handle subsequent memory demands, the second VM will
have to relinquish its resources. We believe that free memory
balancing should be explicit: a guest with low memory utilisation
should have its allocation decreased. Several systems that perform
this resource management automatically have been
1 How are duplicates detected?
The objective of any memory sharing mechanism is to reuse the
reclaimed pages in order to pay for the cost of running the sharing
machinery. A common approach is to add the extra memory to a global
pool, which can be used to create additional
VMs . However, we believe that only the VMs that
participate in sharing should reap the benefits of additional
memory. This creates an incentive for VMs to share memory, and
prevents malicious VMs from negatively affecting the performance of
other VMs on the same host. Therefore, Satori distributes reclaimed memory in proportion to the
amount of memory that each VM shares.
2 How are memory savings distributed?
When Satori identifies duplicates of the same page, it will
reclaim redundant copies. In the common case of , our
policy awards each of the contributing VMs with an entitlement of
pages--or, more generally, pages--for
each shared page (Figure 2). For
each page of physical memory, , we define
the sharing rank of that page. For VM , which uses the set
of pages , the total sharing entitlement, , is
calculated as follows:
Sharing entitlement calculation
Satori interrogates the sharing entitlements for each VM every second, and makes
the appropriate amount of memory available to the VMs.
The sharing rank of a particular page will not necessarily remain
constant through the lifetime of the sharing, since additional
duplicates may be found, and existing duplicates may be
removed. Therefore, the sharing entitlement arising from that page may
change. Consider what happens when a new duplicate is discovered for
an already -way shared page. The VM that provided the new duplicate
will receive an entitlement of pages, and the owners
of the existing duplicates will see their entitlement increase by
for each copy they
own. Similarly, the entitlements must be adjusted when a shared page
In Satori, guests claim their sharing entitlement using memory
balloons . When the entitlement
increases, the balloon deflates and releases additional pages to the
guest kernel. In our implementation we set up the guests to always
claim memory as soon as it becomes available. However, guests can
elect to use more complex policies. For example a guest may refrain
from using its entitlement if it experiences low memory demand, or
expects its shared pages to be short-lived. We have explicitly avoided
using host paging to deal with fluctuating memory
allocations. As a result, our implementation is simpler, and we have
avoided the well-known problems associated with host paging.
However, without host paging, we have to
guarantee that the hypervisor can recover memory from the guests when
it needs to create private copies of previously-shared pages. In the
next subsection, we introduce the repayment FIFO, which
addresses this issue.
If two or more VMs share the same copy of a page, and one VM attempts
to write to it, the VMM makes a private copy of the page. Where does
the VMM get memory for this copy?
Satori obtains this memory from a guest-maintained repayment
FIFO, which contains a list of pages the guest is willing to give
up without prior notification. The size of a VM's repayment
FIFO must be greater than or equal to its sharing entitlement. Our
approach has three major advantages: (a) the hypervisor can obtain
pages quickly, as there is no synchronous involvement with the guest,
(b) there is no need for host paging, and (c) there is no risk that
guest will be unable to relinquish resources due to double
copy-on-write faults (i.e. a fault in the copy-on-write fault
Pages in the repayment FIFO must not contain any irreplaceable
information, because the guest will not have a chance to save their
contents before the hypervisor reclaims them. Memory management
subsystems already maintain book-keeping information about each page,
which makes it possible to nominate such volatile pages without
In Satori the hypervisor uses sharing entitlements to determine the VM from
which to reclaim memory. It does so by inspecting how much memory each VM drew
from the sharing mechanism, in comparison to its current sharing entitlement.
Since the sum of sharing entitlements is guaranteed to be smaller or equal to
the number of removed duplicate pages, there will always be at least
one VM with a
negative memory balance (i.e. the VM drew more than its entitlement). Note that only the
VMs which are involved in the broken sharing will be affected. This is essential
to maintain performance isolation, as a malicious VM will be unable to affect any
VMs with which it does not share memory.
A special case of broken sharing is when a page is reallocated for
another purpose. For example, a guest may decide to evict a shared
page from the page cache, scrub its content and reallocate it. In a
copy-on-write system, the scrubber would cause a page fault when it
begins to scrub the page, and the VMM would wastefully copy the old
contents to produce a private version of the page. We use a scheme
called no-copy-on-write, which informs the VMM that a page is
being reallocated, and instead allocates a zero page (from a
pre-scrubbed pool) for the private version.
To the best of our knowledge, Satori is the first system to
address a covert channel created by memory sharing. An
attacker can infer the contents of a page in another guest, by inducing
sharing with that page and measuring the amount of time it takes to
complete a write. (If a page has been shared, the processing of a
copy-on-write fault will measurably increase the write latency.)
For example, we might want to protect the identity of server processes
running in a guest, because security vulnerabilities might later be
found in them. We allow guests to protect sensitive data by specifying
which pages should never be shared. Any attempts to share with these
pages will be ignored by the hypervisor.
3 What if sharing is broken?
We implemented Satori for Xen version and Linux version
in lines of code ( in the Xen hypervisor, in the
Xen tools and in Linux). We chose Xen because it has extensive
support for paravirtualised guests . In this
section, we describe how we implemented the design decisions from
Our changes can be broken down into three main categories. We first
modified the Xen hypervisor, in order to add support for sharing pages
between VMs (§ 5.1). Next, we added support for
sharing-aware block devices to the Xen control tools
(§ 5.2). Finally, we enlightened
the guest operating system, so that it can take advantage of
additional memory and repay that memory when necessary
The majority of our changes were contained in the hypervisor. First of
all, the upstream version of Xen does not support transparent page
sharing between VMs, so it was necessary to modify the memory
management subsystem. Once this support was in place, we added a
hypercall interface that the control tools use to inform the
hypervisor that pages may potentially be shared. Finally, we modified
the page fault handler to deal with instances of broken sharing.
In § 2.2, we explained that each VM has a contiguous,
zero-based pseudo-physical address space, and a P2M mapping for converting
pseudo-physical addresses to machine addresses. To support transparent page
sharing, it is necessary to allow multiple pseudo-physical pages to map to a
single frame of machine memory. Furthermore, the machine frame that backs a
given pseudo-physical page may change due to sharing. Therefore, it is simplest
to use shadow page tables in the guest VMs. However, regular paravirtualised
guests in Xen do not use shadow page tables, so we ported this feature from the
code which supports fully-virtualised guests. In addition, we had to modify the
reference counting mechanism used in Xen to keep track of page owners. In Xen
each page has a single owner, so we added a synthetic ``sharing domain'' which
owns all shared pages.
As described in § 5.3, we maintain information
about the state of each (pseudo-)physical page in each guest. Both the
guest and the hypervisor may update this information, so it is held in
a structure that is shared between the hypervisor and the
guest. The hypervisor uses this structure to select which page should
be used to satisfy a copy-on-write fault (either a page from the
repayment FIFO, or, in the no-copy-on-write case, a zero-page).
We export the sharing functionality to the guest through the hypercall
interface. We add three new hypercalls, named
share_mfns, mark_ro and get_ro_ref.
The share_mfns hypercall takes two machine frame numbers
(MFNs)--a source MFN and a client MFN--and informs the hypervisor that all
pseudo-physical pages backed by the client frame should now use the source
frame. The hypercall works as follows:
1 Hypervisor modifications
Note that this hypercall is not guaranteed to succeed. For example,
after the duplicate detector notices that two pages are the same, but
before they are marked read only, a guest might change the contents of
one of the pages. Therefore, the hypercall may fail, but there is no
risk that the contents of memory will be incorrect: the source and
client frame will continue to be used as before.
For copy-on-write disks, we want to make an early decision about
whether or not physical I/O will be required. Therefore, we use the
mark_ro hypercall to enforce read-only status on all pages that are
read from the read-only substrate. (Technically, we make a page read-only by
treating it as 1-way shared; if the guest writes to it, the sharing is simply
broken by marking the page as writable and changing the owner to the guest.) The
complementary get_ro_ref hypercall ensures that the contents of the
frame have not been changed (i.e. that the MFN is still read-only), and
increments the sharing reference count to prevent it from being discarded. We
describe the copy-on-write disk support in § 5.2.
The final hypervisor component required for page sharing is a modified
page fault handler. We added two new types of page fault, which Xen
must handle differently. The first is a straightforward
copy-on-write fault, which is triggered when a guest attempts
to write to a shared page. In this case, the handler recalculates the
sharing entitlements for the affected guests, and reclaims a page from
one of the guests that now has claimed more memory than its
entitlement. The handler removes this page from the appropriate
guest's repayment FIFO and copies in the contents of the faulting
page. We also add a discard fault, which arises when a guest
attempts to access a previously-volatile page that the VMM has
reclaimed. If so, the handler injects this fault into the guest, as
described in § 5.3.
- Mark the source and client frame as read-only, if they are not
- Compare the contents of the source and client frame. If they are
not equal, return an error.
- Remove all mappings to the client MFN from the shadow page
- Update the relevant P2M mappings to indicate
that the source frame should be used in place of the client frame.
- Free the client frame for use by the guest VMs.
We implemented duplicate detection using sharing-aware block
devices. Xen provides a high-performance, flexible interface for
block I/O using split devices. The guest contains a
front-end driver, which presents itself to the guest OS as a
regular block device, while the control VM hosts a corresponding
back-end driver. Previous work has shown how the back-end is a
suitable interposition point for various
applications , in particular for creating a
distributed storage system . We use the existing
block-tap architecture to add duplicate detection.
The key steps in a block-tap read request are as follows:
2 Sharing-aware block devices
Since tapdisk is implemented as a user-space process and provides
access to I/O data, it is simple to add custom block-handling code at
this point. Satori modifies the tapdisk read path in order to record
information about what data is loaded into which locations in the
guests' memory. We developed two versions of duplicate detection:
content-based sharing, and copy-on-write disk sharing.
For content-based sharing,
we hash the contents of each block as it is read from disk. We use the
hash as the key in a hashtable, which stores mappings from hash values
to machine frame numbers (MFNs). First, we look for a match in this
table, and, if this is successful, the resulting MFN is a candidate
for sharing with the I/O buffer. Note that the MFN is merely a hint:
the contents of that frame could have changed, but since we have
already loaded the data into the I/O buffer, it is acceptable for the
sharing attempt to fail. If the hash is not present in the hashtable,
we invalidate any previous entry that maps to the I/O buffer's MFN,
and store the new hash-to-MFN mapping.
- The front-end (in the guest) issues a read request to the
back-end through the inter-VM device channel, by providing
a block number and an I/O buffer.
- The back-end maps the I/O buffer into a user-space control
tool, called tapdisk.
- tapdisk performs device-specific processing for the
given block number, and returns control to the back-end driver.
- The back-end unmaps the I/O buffer and notifies the front-end
For copy-on-write disk sharing, the process is slightly different (see
Figure 3). The first time a block is
read from the substrate disk, Satori invokes the mark_ro
hypercall on that page, and stores a mapping from the block
number to the I/O buffer MFN. (If the guest subsequently writes to
the page before it is shared, the read-only status is removed.) On
subsequent reads, Satori consults the block number-to-MFN mapping to
see if the block is already cached in memory. If it is, Satori invokes
the get_ro_ref hypercall on the MFN, which, if it succeeds,
ensures that the subsequent call to share_mfns will be
successful. If the block number is not found in the mapping table or
if get_ro_ref fails, Satori must request physical disk I/O
to read the appropriate block. At this point, a second look-up could
be used to detect content-based sharing opportunities.
The Xen architecture places each virtual block device in a
separate tapdisk process, to simplify management and
improve fault isolation. However, we need to share information between
devices, so it was necessary to add a single additional process, called
spcctrl (Shared Page Cache
ConTRoLler), which hosts the mappings
between content hashes or block numbers, and MFNs. The
tapdisk processes communicate with spcctrl using pipes, and our
current implementation of spcctrl is single-threaded.
mark_ro and get_ro_ref usage for copy-on-write disks.
In Satori, we have used enlightenments to obtain OS-level information
about guest pages. These extend the existing paravirtualised Linux
guests, which Xen already supports .
For Satori, the most important enlightenment is adding the repayment FIFO
to the guest kernel. Recall that the repayment FIFO is a list of
volatile pages, i.e. physical pages that the operating system
is willing to relinquish at any time (in particular, when sharing is
broken and a new frame is needed for a private copy). Since the guest
must relinquish these pages without warning, it is essential that
their contents can be reconstructed from another source. Hence an
obvious source of volatile pages is the set of clean pages in the page
cache. We paravirtualised the Linux page cache to provide page
hints about volatile pages.
We based our implementation of volatile pages on earlier work on
Collaborative Memory Management
(CMM) . CMM is, in essence, a
memory controller which relies on page states (especially page
volatility) to dynamically adjust the available (machine) memory for
each guest. CMM is implemented for the IBM zSeries z/VM hypervisor,
but the majority of the code is architecture-independent, as it deals
with page-state transitions in the Linux page and swap caches. We
built on CMM's page hinting by adding support for the x86 architecture
and the Xen hypervisor.
The major difference between x86 and zSeries (s390) in the context of
volatile pages is the handling of dirtying. The x86 architecture
maintains dirty bits for virtual pages in the PTE, whereas the s390
architecture maintains a dirty bit for each machine page. Since a
given page can only become volatile if it is not dirty, we implemented
a machine-page-level dirty bit in software for the x86 architecture. Our
approach is more conservative than is strictly necessary, because we
consider the existence of any writable mapping to dirty the page, even
if there was no actual write.
Satori uses a shared structure between Xen and each guest to store and
modify page states (as discussed in § 5.1). The
page states read from this structure are used in the guest page fault
handler to distinguish between ``regular'' and discard
faults. On a discard fault, Linux uses reverse mappings to remove all
references to the discarded page, effectively removing the page from its
respective cache (page or swap).
We also use the instrumentation in the page allocator, already present
in order to drive page state transitions, to support the
no-copy-on-write policy. Whenever a page is reallocated, we update the
shared page state structure to reflect this. On a write fault to a
shared page, Xen checks to see whether the page has been reallocated,
and, if so, provides a page from its zero page cache.
In addition, we have added support to guests for specifying that some
pages must not be shared (to avoid the secret-stealing attack
described in § 4.3).
At present, we allow the
guest to specify that a set of pseudo-physical pages must never be
shared (i.e. all calls to share_mfns or
get_ro_ref will fail).
3 Guest enlightenments
To characterise Satori's performance, we have conducted an evaluation in three
parts. First, we have profiled the opportunities for page sharing
under different workloads (§ 6.1). In contrast
with previous work, we specifically consider the duration of
each sharing opportunity, as this is crucial to the utility of page
sharing. We then measure the effectiveness of Satori, and show that it
is capable of quickly detecting a large amount of sharing
(§ 6.2). Finally, we measure the effect that
Satori has on performance, in terms of the benefit when
sharing is enabled, and the overhead on I/O operations
For our tests we used two Dell PowerEdge 1425 servers each equipped with two
GHz Xeon CPUs, GB of RAM and an GB Seagate SATA disk. VMs ran
Ubuntu Linux 8.04 in all cases, except for two experiments, for which we state
the OS version explicitly.
In the following subsections, we make repeated reference to several
workloads, which we abbreviate as follows:
|| ||Vanilla Linux kernel build with 256 MB of
|| ||As KBUILD-256, with MB.
|| ||httperf benchmark  run against Apache web-server with 512 MB of
memory, serving randomly generated static webpages.
|| ||RUBiS web auction application with MB, serving
requests generated by the default client workload generator .
The major difference between Satori and contemporary page sharing
schemes is that it can share many identical pages as soon as they are
populated. In this subsection, we show that a substantial proportion
of sharing is short-lived. Therefore, Satori is much more likely to
exploit this sharing than schemes that rely on periodically
scanning physical memory, looking for identical page
To analyse the sharing opportunities, we ran each of the KBUILD-256, KBUILD-512,
HTTPERF and RUBIS workloads in two virtual machines for minutes,
and took a memory dump every seconds.
1 Sharing opportunities
Breakdown of sharing opportunities by rank (excluding zero-pages).
Table 1 shows the number of pages that can be
freed using page sharing, for each rank. (In a sharing of rank ,
identical pages map to a single physical page.) The figures are an aggregate,
based on the total of 60 memory dumps sampled from pairs of VMs running the
KBUILD-512, HTTPERF and RUBIS workloads. Note that most sharing opportunities have
rank : i.e. two identical pages exist and can be combined into a single
Sharing opportunities during the execution of workloads KBUILD-256 and KBUILD-512.
Duration of page sharing opportunities for kernel compilation workloads.
(a) and (b) show non-zero pages, (c) and (d) zero pages.
The exploded sectors show sharings left at the end of the experiment.
Figure 4 compares the number
of unique shared pages during the KBUILD-256 and
KBUILD-512 workloads. (By considering only unique shared pages, we
underestimate the amount of savings for pages with rank . Table 1 demonstrates that the
majority of shareable pages have rank , except zero pages, which we
address separately below.) We have divided the sharing opportunities
into four duration ranges. The figures clearly show that a substantial amount of sharing
is short-lived, especially in a more memory-constrained setup
(KBUILD-256). Also, the amount of sharing for the KBUILD-512 workload is
approximately twice as much as that for KBUILD-256, because of less
contention in the page cache. Finally, the kernel build process
completes minutes sooner with MB of memory: this makes the
benefits of additional memory clear.
Breakdown of Satori hypercalls during HTTPERF workload
Figure 5 separately categorises shareable non-zero
pages and zero pages into the same duration ranges as
Figure 4. It should be noted
that the number of sharing opportunities arising from zero pages
(Figures 5(c) and 5(d)) is
approximately times greater than from non-zero pages
and 5(b)). However, more than of
zero-page sharing opportunities exist for less than five minutes. This
supports our argument that the benefits of zero-page
sharing are illusory.
In § 4.1, we stated
that, on average, it will take minutes for VMware ESX Server to
detect a duplicate page using its default page scanning options. We
ran the following experiment to validate this claim. Two VMs ran a
process which read the same MB, randomly-generated file into
memory, and Figure 6 shows the number of
shared pages as time progresses. The lower curve, representing the
default settings, shows that half of the file contents are shared
after minutes, which is close to our predicted value; the
acceleration is likely due to undocumented optimisations in VMware
ESX Server. The higher curve shows the results of the same experiment
when using the most aggressive scanning options. Using the same
analysis, we would expect a duplicate on average to be detected after
minutes. In our experiment, half the pages were detected after
almost minutes, and we suspect that this is a result of the
aggressive settings causing the page hint cache to be flushed more
In the next set of experiments, we measured the amount of sharing that Satori
achieved using sharing-aware block devices. We also examined how the
surplus memory was distributed between individual virtual machines.
The first experiment used two pairs of virtual machines. Two VMs each
ran the HTTPERF-256 workload, i.e. the HTTPERF workload with MB of
memory (rather than MB). Because the aggregate amount of memory
was insufficient to cache the entire data set in memory, the number of
shareable pages varied as data was loaded into and evicted from each
VM's page cache. The other two VMs each ran the KBUILD-512 workload;
however they used Debian Linux rather than Ubuntu.
Figure 7 shows that the sharing
entitlements for the VMs running KBUILD-512 are unaffected by the highly
variable amount of sharing between the two HTTPERF workloads. Also,
because we used different OSes for each pair of VMs, the sharing
entitlements achieved before the workloads started (5 to 6 minutes
after the measurements began) differ by about 30%.
Next, we ran two instances of a workload in separate VMs for
minutes, and repeated the experiment for the KBUILD-256, KBUILD-512, HTTPERF and
RUBIS workloads. We ran these experiments under Satori and measured
the number of shared pages, and compared these to memory dumps using
the same methodology as described in § 6.1.
Figure 8 summarises the amount of sharing
that Satori achieves for each workload. Satori performs best with the
HTTPERF workload, shown in Figure 8(c). In
this case, it achieves of the total sharing opportunities,
which is to be expected since HTTPERF involves serving a large volume
of static content over HTTP, and the majority of the data is read
straight from disk. The RUBIS workload performs similarly, with Satori
achieving of the total. The kernel compilation
workloads, KBUILD-256 and KBUILD-512, perform less well. KBUILD-512 achieves about
of the total sharing opportunities until the very end of the
build, when the kernel image is assembled from individual object
files. KBUILD-256 is more memory-constrained, which forces the OS to flush
dirty (non-shareable) caches.
2 Satori effectiveness
Sharing as time progresses for default and aggressive
scanning settings in VMware ESX Server.
Finally, we ran two experiments which evaluated Satori in a more
heterogeneous environment. In the first experiment, two VMs running
the same version of Ubuntu Linux performed the HTTPERF and
RUBIS workloads. In this setup Satori was able to exploit over
of the total sharing opportunities. (The remaining was
mostly due to the identical kernel images, which the current version
of Satori does not detect.). In the second experiment, we used the
same workloads with different guest OSs (Ubuntu and Debian
respectively). In this setup, MB of sharing was
theoretically possible, and only because the two distributions use an
identical kernel. In this case, Satori could only achieve
approximately MB of savings ( of the total).
Sharing entitlements for two KBUILD-512 and two HTTPERF-256
workloads executing simultaneously.
Although Satori achieves varying results in terms of memory savings,
recall that these results come solely from using our enlightened block
device. These results show that we can exploit up to
(for HTTPERF) of the total sharing opportunities
through this method alone. The alternative approach, which involves
scanning memory, incurs a constant overhead at run-time, and must be
rate-limited to prevent performance
degradation . The Difference Engine
exhibits an overhead of up to on some macrobenchmarks, though this
includes the overhead of page compression and sub-page-level
patching . Satori provides a flexible interface for
adding other sharing policies: we are developing a tool that
systemically identifies the source(s) of other sharing
opportunities. We hope that this will lead to additional
enlightenments that improve Satori's coverage.
In § 4.3, we described an attack on memory
sharing that allows a VM to identify sensitive data in another
guest. On VMware ESX Server, we were able to determine the precise
version of sshd running in another guest, by loading a page
from each of common distribution-supplied
sshd binaries into memory, and periodically measuring the
write latency to these pages. (On our hardware, we observed a
-times increase for the matching page.) In Satori, we were able to
protect the entire sshd address space, and, as a result, this
Amount of sharing achieved by Satori for each of the four
main workloads (no zero-pages)
We initially stated that memory sharing is desirable because it can
improve the overall performance of VMs running on the same physical
machine. In this subsection, we investigate the performance
characteristics of Satori under various workloads. First, we measure
negative impact: Satori introduces new operations for sharing
memory, and these incur a measurable cost. We measure the cost of each
sharing-related hypercall, and the overall effect on disk
I/O. However, we then show that, for realistic macrobenchmarks, the
overhead is insignificant, and the additional memory can improve
To measure the cost of individual sharing-related operations, we
instrumented the Xen hypervisor to record the number and duration of
each hypercall. Table 2 shows the results
for a -minute HTTPERF
workload. The first thing to note is that Satori-related operations account for less
than seconds of the -minute benchmark.
Of the individual operations,
mark_ro is the most expensive, as it must occasionally
perform a brute-force search of the shadow page tables for all
mappings of the page to be marked read-only. We could optimise
performance in this case by making the guest VM exchange
back-reference information with the hypervisor, but the overall
improvement would be negligible.
Satori detects sharing by monitoring block-device reads, and therefore
the majority of its overhead is felt when reading data from disk. In
order to measure this overhead, and stress-test our implementation, we ran the Bonnie filesystem benchmark in a guest VM against a
sharing-aware block device. Table 3 shows a breakdown
of read bandwidths. We ran the benchmark in four
configurations, and repeated each experiment five times. In the
baseline configuration, we disabled all Satori mechanisms. In successive
configurations, we enabled content hashing, IPC with spcctrl,
and finally hash lookup, in order to isolate the performance impact of
Table 3 reports bandwidth figures for
reads using getc(), and ``intelligent reads'', which use a
block size of bytes.
3 Performance impact
Results of the Bonnie filesystem benchmark on Satori
||Read bandwidth (MB/s)
|Hashing + IPC
The first thing to note is that Bonnie performs sequential
reads on a MB file, so the effect of any computation on the I/O
path is amplified. (The impact of Satori on random reads is
negligible.) Therefore, the overhead for chunked reads with
Satori fully enabled is a worst-case figure, and is to be
expected. With a realistic workload, the I/O cost is likely to
dominate. Nevertheless, it is instructive to consider the individual
While we are continuing to improve Satori's performance, and eliminate
these bottlenecks, we have only encountered the above issues when
running Bonnie. For example, we ran a stripped-down kernel
compilation in a single VM, which took an average of
seconds with Satori disabled, and seconds with Satori
fully enabled. Since the standard deviation over five runs was
seconds, it is clear that the overhead is statistically insignificant.
In this experiment, the workload ran in isolation, and there were no
benefits from sharing. As we will see next, the advantage of having
additional memory can improve performance for many workloads.
We first ran an experiment to illustrate the benefit of memory sharing
between VMs that share a copy-on-write disk. We ran a workload
that read the contents of a copy-on-write disk into memory in a
pseudorandom order. Five seconds later (while the first read was
ongoing), we started the same workload, reading from the same disk in
the same sequence, in another
VM. Figure 9 shows the progress that both
VMs achieved as a proportional gradient. VM1 reads at a consistent rate of
MB/s. When the workload commences in VM2, its initial
rate is MB/s, as the data that it reads can
be provided by sharing memory with the page cache in
VM1. After seconds, VM2 has read all
of the data held by VM1, and the two VMs become mutually synchronised,
at the original rate, which is limited by the disk access
time. Although this example is artificial, it shows Satori's
effectiveness at exploiting page cache sharing for copy-on-write
disks. Many recent cloud computing systems, such as Amazon's
EC2 , encourage the use of standard machine image
files, which are natural candidates for a copy-on-write
implementation. Satori would be particularly effective in this case.
- The overhead involved in hashing is relatively constant and less
- IPC with the spcctrl process is costly. The present
implementation uses UNIX pipes to communicate with spcctrl, which
involves two additional context switches per read request. We plan to redesign
this component to store the hashtable in a shared memory segment.
- The relative overhead of fully-enabled Satori is worse in the
chunked read case, because less time is being wasted in making
repeated system calls in the guest VM.
Finally, we ran the HTTPERF workload in two VMs as a macrobenchmark, to
discover how well Satori exploits the extra memory that is made
available through sharing. We compare Satori to VMware ESX Server--the
leading commercial hypervisor--which uses the
techniques described by Waldspurger to achieve page sharing and memory
Copy-on-write disk read rates
Figure 10 shows how the aggregate
HTTPERF response rate changes over time for Satori and VMware
(with and without VMware Tools). The performance of Satori can be
divided into two phases. First, it achieves approximately
responses per second while the cache is being loaded, which takes
approximately seconds. The response rate then jumps to between
and responses per second as all subsequent requests can be
satisfied from caches. In order to maintain these response rates, the
VMs use their sharing entitlements to increase their page cache
sizes. The physical memory available to each VM grows to over MB
over the first seconds of the experiment.
The results for VMware are interesting. We note first that it was
necessary to install the VMware Tools (which include a balloon driver)
in order to achieve performance that was comparable to Satori. Without
the VMware Tools, the VMM begins paging after approximately
seconds, and throughput drops almost to zero. Once host paging starts,
the throughput only recovers occasionally, and never to more than
responses per second.
With the VMware
Tools installed, we observed that balloon permanently limited
each VMs physical memory allocation to MB. Therefore, the VMs were able to
make progress without host paging, but the data set did not fit in the
cache, and the response rate remained at around responses per
second. VMware was unable to establish sufficient sharing because the
lifetime of a page in either page cache was usually too short for the memory
scanner to find it.
Aggregate HTTPERF response rates for the two VMs running on
Satori, VMware, and VMware with VMware
We described Satori, which employs enlightenments to improve the
effectiveness and efficiency of page sharing in virtualised
environments. We have identified several cases where the traditional page
sharing approach (i.e. periodic memory scanning) does not discover or
exploit opportunities for sharing. We have shown that, by using
information from the guest VMs, and making small modifications to the
operating systems, it is possible to discover a large fraction of the
sharing opportunities with insignificant overhead.
Our implementation has concentrated on sharing-aware block
devices. In the future we intend to add other enlightened page
sharing mechanisms--such as long-lived zero-page detection,
page-table sharing and kernel text sharing--which will improve
Satori's sharing discovery rate. We also intend to investigate the
application of our technique to nearly-identical
We wish to thank members of the Systems Research Group at the University of
Cambridge for the many fruitful discussions that inspired this work. We also
wish to thank our shepherd, Geoffrey Voelker, and the anonymous reviewers for
their insightful comments and suggestions that improved this paper.
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