Computer Science Division
Department of Electrical Engineering and Computer Science
University of California, Berkeley
Berkeley, California 94720
This paper describes a modified version of the UNIX operating system
that supports virtual memory through demand paging.
The particular implementation being described here is specific
to the VAX†-11/780 computer system
although most of the design decisions have wider applicability.
この論文は、 デマンド・ページングによって仮想記憶をサポートする UNIX オペレーティング・システムの修正版について記述します。 ここに記述されている実装方法は特に VAX†-11/780 コンピューター・システムに限定されるものですが、 設計上の決定の大部分はより広い範囲に適用することが可能です。
The modified system creates a large virtual address space for user programs
while supporting the same user level interface as UNIX.
The few new system calls that have been introduced are
primarily aimed for performance enhancement.
The paging system implements a variant of the global CLOCK
replacement policy (an approximation of the global
least recently used
with a working-set-like mechanism for the control of multiprogramming level.
修正済のシステムは、 UNIX と同じユーザ・レベル・インターフェースを支援するだけでなく、 ユーザ・プログラムのための大きな仮想のアドレス空間を作成します。 新たに導入された幾つかのシステムコールは、 主にパフォーマンスの向上を目的としたものです。 ページ・システムには マルチプログラミング・レベルでの制御のための working-set-like mechanism と共に (the global least recently used algorithm の近似である) the global CLOCK replacement policy を修正して実装しました。
Measurement results indicate that the lack of
in the VAX memory-management hardware can be overcome at relatively
little expense through software detection.
Also included are measurement results comparing
the virtual system performance
to the swap-based system performance
under a script-driven load.
測定結果は、 ソフトウェアによる検知方法により 比較的わずかな費用で VAX メモリ管理ハードウェア中の参照ビットの不足を 克服することができることを示します。 さらに、 測定結果には スクリプトで駆動される負荷を使って、 スワップ・ベースのシステムのパフォーマンスと 仮想システムのパフォーマンスの比較が含まれています。
Keywords and phrases: UNIX, virtual memory, paging, swapping, operating systems, performance evaluation, VAX.
† UNIX and UNIX/32V are Trademarks of Bell Laboratories
† Work supported by the National Science Foundation under grants MCS 7807291, MCS 7824618, MCS 7407644-A03 and by an IBM Graduate Fellowship to the second author.
† VAX and PDP are trademarks of Digital Equipment Corporation.
The most significant architectural enhancement
that the VAX-11/780 provides
over its predecessor, the PDP-11, is the very large address
space made available to user programs.
The fundamental task of transporting UNIX to this new hardware
was accomplished by Bell Laboratories at Holmdel.
In addition to the portability directed changes,
the memory-management mechanism of the base system
was modified to make partial use of the new hardware.
through these changes,
processes could be
scatter loaded into memory
thus avoiding main-memory fragmentation,
and swapped in and out of memory partially.
A process, however, still had to be fully loaded in order to execute.
While no longer limited by the 16 bit address space of the PDP-11,
the per-process address space could grow
only as large as the physical memory available to user processes.
This system, which constituted a prerelease of UNIX/32V,
was adopted as the basis for virtual memory extensions.
VAX-11/780 が、 その先行機種である PDP-11 に対して行われた最も重要なアーキテクチュア上の拡張は、 ユーザ・プログラムに利用可能な非常に大きなアドレス空間です。 この新しいハードウェアに UNIX を移植する基本的な作業は、 Bell Laboratories at Holmdel によって行われました。 ポータビリティに関わる変更に加えて、 ベース・システムのメモリ管理メカニズムは 新しいハードウェアを部分的に利用するために修正されました。 これらの変更によって、 特にプロセスはメモリ上に 離散的に展開する ことができるようになりました。 その結果、メイン・メモリのフラグメンテーションは回避され、 部分的なスワップ・イン／スワップ・アウトが可能になりました。 しかしながら、 実行をするためにはプロセスを完全に展開しなければなりませんでした。 もはや PDP-11 の１６ビットのアドレス空間に制限されることがないだけではなく、 プロセスごとのアドレス空間は ユーザープロセスに利用可能な物理メモリーと 同じサイズまで大きくすることができます。 UNIX/32Vのプリ・リリースを構成するこのシステムは、 仮想記憶のための拡張のベースとして採用されました。
The virtual memory effort was motivated
by several factors in our research environment:
The reader should be familiar with the standard UNIX
system as described in [RITC 74]
and the virtual memory concept in general [DENN 70].
In the next section, we briefly describe the memory-management hardware
as it exists in the VAX-11/780 to support virtual memory [DEC 78].
The following sections detail the new kernel operations
including new system calls followed by various measurement results.
この論文の読者は、 [RITC 74] で記述されるような 標準的な UNIX システム、 および仮想記憶の概念に関する一般的な知識 [DENN 70] に精通しているべきです。 次のセクションでは、 仮想記憶をサポートするために VAX-11/780 に搭載されているメモリ管理ハードウェア [DEC 78] について簡潔に説明します。 続くセクションでは、 様々な測定結果を示して、 新しいシステムコールを含む、 新しいカーネルの操作の詳細を述べます。
The VAX-11/780 memory-management hardware supports
a two level mapping mechanism to perform the address translation task.
The first level page tables reside in system virtual address space
and map user page tables. These tables in turn, map the
user virtual address space which consists of 512 byte pages.
The 32 bit virtual address space of the VAX-11/780
is divided into four equal sized blocks.
VAX-11/780 のメモリ管理ハードウェアは、 アドレス変換タスクを実行するために ２レベルのマッピング・メカニズムをサポートします。 第１レベルのページ・テーブルには、 システム仮想のアドレス空間とユーザ・ページ・テーブルのマップが存在します。 これらのテーブルは、 ５１２バイトのページから構成される、 ユーザの仮想のアドレス空間を順番にマップします。 VAX-11/780 の３２ビットの仮想のアドレス空間は、 同じ大きさの４つのブロックに分割されます。
Two of these blocks, known as the P0 and P1 regions,
constitute the two per-process segments.
The third block, known as the system region,
contains the shared kernel virtual address space
while the fourth region is not supported by the current hardware.
The P0 segment starts at virtual address 0 and can grow
toward higher addresses.
The P1 segment on the other hand, starts at the top of user
virtual address space and grows toward lower addresses.
Both segments are described by two per-process (base, length) register pairs.
このブロックのうちの２つは、 P0/P1 のリージョンとして知られ、 プロセスごとの２つのセグメントを構成します。 第３のブロックはシステム・リージョンとして知られ、 共有されるカーネル仮想アドレス空間が含まれます。 第４のリージョンは現在のハードウェアではサポートされていません。 P0 セグメントは仮想アドレスの０からスタートし、 より高いアドレスの方へ成長することができます。 一方、 P1 セグメントはユーザの一番上に仮想のアドレス空間から始まり、 下位アドレスの方へ成長します。 両方のセグメントは、 プロセス毎に ２組のレジスタ・ペア (base, length) によって記述されます。
A page table entry consists of
four bytes of mapping and protection information.
Attempting a translation through a page table entry
that has the valid bit off results in a
Translation Not Valid Fault (i.e., a page fault).
Whereas most architectures that support virtual memory
provide a per-page Reference Bit
that is automatically set by the hardware
when the corresponding page is referenced
to be examined and/or reset
by the page replacement algorithm,
has no such mechanism.
In order to overcome this deficiency, we distinguish the three states
that a page of virtual address space can be in:
ページ・テーブル・エントリーは、 マッッピングと保護に関する情報を示す４バイトから構成されます。 valid ビットをオフにしたページ・テーブル・エントリーによる、 変換の試みは Translation Not Valid Fault という結果になります。(つまりページフォルト) 仮想記憶をサポートするほとんどのアーキテクチャでは、 対応するページが参照された時に ハードウェアにより自動的にセットされるページ毎の Reference Bit が提供されます。 これはページ置換アルゴリズムによって 参照あるいはリセットするために使われますが、 VAX-11/780 はそのようなメカニズムを持っていません。 この欠陥を克服するために、 私たちは仮想アドレス空間のページが取り得る３つの状態を識別します:
This scheme in effect
allows us to detect and record references to pages using software.
We discuss the cost and effectiveness of the method in §7.2.
このスキームは事実上、 ソフトウェアを使用して、 ページへの参照を検知し記録することを可能にします。 私たちは §7.2. で、 この方法のコストと有効性について議論します。
the notion of a process and a computer execution environment
are intimately related [THOM 78].
In fact, a process is the execution of this environment
which consists of the process virtual address space state,
general register contents, open files, current directory, etc.
The state of this pseudo computer
is comprised of the contents of four segments.
The first three contain the process virtual address space,
while the fourth segment describes
the system maintained state information.
UNIX では、 コンピューターの実行環境と process という概念が密接に関連づけられます。 実際、 process とは プロセスのアドレス空間の状態、 一般的なレジスタ内容、 オープンされているファイル、 カレント・ディレクトリ などから構成される環境を実行することを意味します。 この疑似コンピュータの状態は ４つのセグメントの内容で成立しています。 ４番目のセグメントには システムが維持している状態情報が記述されているのに対し、 最初の３つにはプロセスの仮想アドレス空間が含まれます。
The process virtual address space
consists of a read only program text segment
that is shared amongst all processes
that are currently executing the same program,
as well as private writable data and stack segments.
Within the limited segmentation capability of the
these three segments are mapped
such that the program text is in the P0 region beginning
at virtual address 0
with the data immediately after it starting at the next page boundary.
The stack segment is mapped into the P1 region starting
at the highest virtual address.
While the text segment has a static size,
the data segment can be grown or shrunk through system calls
and the stack segment is grown automatically
by the kernel upon the detection of segmentation faults.
仮想アドレス空間は、 書き込み可能なデータおよびスタック・セグメント、 および同じプログラムを実行している時に全てのプロセスの間で共有される 読み出し専用のプログラム・テキスト・セグメントから構成されます。 VAX-11/780 のセグメントの機能的な制限により、 これらの３つのセグメントは次のようにマッピングされます。 プログラム・テキストは仮想アドレス０から始まる位置に、 データはプログラム・テキストの次のページ境界から、 いずれも P0 領域にマッピングされます。 スタック・セグメントは 最も高い仮想アドレスからスタートする P1 領域にマッピングされます。 テキスト・セグメントが静的なサイズを持っているのに対し、 データ・セグメントはシステムコールによって拡大／縮小する可能性があります。 また、スタック・セグメントは、 カーネルがセグメンテーション・フォルトを検知すると自動的に拡大します。
The physical structure of these segments in secondary storage
(called an image)
can be organized in various ways.
At one extreme is the physically contiguous organization
described simply by a (base, length) pair.
While appropriate for static segments, such as text,
this organization is too rigid for dynamically growing segments,
like the data and stack segments.
In addition to fragmentation,
segment growth beyond the current allocation
could imply physical movement of the image.
At the other extreme is a fully scattered organization of the image.
While minimizing fragmentation,
this can result in expensive allocation and mapping functions
due to the large number of pages which are present in large images.
２次記憶装置の中のこれらのセグメントの物理的な構造 (imageと呼びます) は様々な方法で構成することができます。 １つの極端な例は、 単に (base, length) の組合せを記述した物理的に連続した構成をとる方法です。 この構成は、 テキストのような静的なセグメントでは適切ですが、 データやスタック・セグメントのように 動的にセグメントが拡大する場合には融通が効きません。 フラグメンテーションが発生する場合、確保されている領域を越えて 拡大するセグメントでは image が物理的に移動することを意味します。 フラグメンテーションを最小限に抑えた場合、 大きな image による大量のページが存在すると、 アロケーションとマッピングの機能は高くつく結果になります。
The image organization chosen for the dynamic segments
represents a compromise between the two extremes.
Each image consists of several scattered chunks.
An exponentially increasing chunk allocation scheme
allows the mapping of very large segments through a small table.
Limiting the maximum size of any chunk helps to prevent extreme fragmentation.
Thus in the current system,
the smallest chunk allocated to a segment is 8K bytes,
and chunk sizes increase by powers of two up to a maximum size of 2M bytes.
動的なセグメントのために選ばれた image 構成は、 ２つの極端な例の妥協を表わします。 各 image はいくつかの散在した chunk から構成されます。 指数関数的に増加する chunk の付けスキームは、 小さなテーブルを介して非常に大きなセグメントのマッピングを許します。 chunk の最大サイズを制限することは、 極端なフラグメンテーションを防ぐことに役立ちます。 したがって現在のシステムでは、 セグメントに割り付けられた最小の chunk は 8K バイトで、 そのサイズは最大 2M バイトまでの２の累乗によって増加します。
We now describe the major new functions performed by the kernel as well as the existing functions whose implementation have been significantly modified. For the purposes of future discussion, we define the following terms:
The most visible of the kernel changes is the existence of a Translation Not Valid fault handler. Given the virtual address that caused the fault, the system checks to see if the page containing the virtual address is in the reclaimable state. This happens when the pageout daemon has swept past a page and made it reclaimable to simulate a reference bit (as described below). If the page is in this state, it can once again be made valid, and the process returns to user mode. Note that if the reclaimed page was in the free list, it is removed and reenters the loop. Since none of the operations involved in reclaiming a page can cause the process to block, reclaiming a page does not involve a processor context switch and reschedule.
If the page cannot be reclaimed (i.e., is not no longer in core), then a page frame is allocated and the disk transfer is initiated from the segment image as dictated by the image mapping.
In reality, more cases must be considered. If the faulting page belongs to a shared text segment, the disk transfer is initiated only if the page is not reclaimable and not intransit, i.e., the pagein operation has not already been initiated by another process that is sharing the text segment. If intransit, the faulting process sleeps to be waken by the process that started the page transfer when it completes. Here we note that the first level page tables for shared text segments are not shared, but rather, each process has its own copy.†
† Sharing all user level page tables of shared segments would require a 64K byte alignment between the text and data segments. This is not enforced by the current loading scheme, so currently these page tables are not shared at all.
Thus, all operations that modify page table entries of shared text segments must insure that all existing copies are updated.
Other types of page faults that require special handling are those where the faulting page is marked as fill on demand. There are three types of demand fill pages:
During system initialization, just before the init process is created, the bootstrapping code creates process 2 which is known as the pageout daemon. It is this process that actually implements the page replacement policy as well as writing back modified pages. The process leaves its normal dormant state upon being waken up due to the memory free list size dropping below an upper threshold.
At this point, the daemon examines the page frame being pointed to by the hand. If the page frame corresponds to a valid page, it is made reclaimable. Otherwise the page was reclaimable, and it is freed, but remains reclaimable until it is removed from the free list and allocated to another purpose. The hand is then incremented and the above steps are repeated until either the free memory is above the upper threshold or the angular velocity of the hand exceeds a bound.
The rate at which the daemon examines page frames increases linearly between the free memory upper threshold and lower threshold (these are tunable system parameters). In a loaded system the hand will be moved around the loop two to three times a minute.
Upon encountering a reclaimable page that has been modified since it was last paged in, the daemon must arrange for it to be written back before the page frame containing it can be freed. To accomplish this, the daemon queues a descriptor of the I/O operation for the paging device driver. Upon completion of the I/O, the interrupt service routine inserts the descriptor to the cleaned list for further processing by the daemon. The daemon periodically empties the cleaned list by freeing the page frames on it that have not been reclaimed in the meantime.
Note that as described above, this pageout process implements a variant of the global CLOCK page replacement algorithm that is known as scheduled sweep [CORB 68, EAST 79].
In UNIX, every process comes into being through a fork system call whereby a copy of the parent process is created and identified as the child. This involves the duplication of the parent's private segments (data and stack) and the system maintained state information (open files, current directory, etc.).
Within a virtual memory environment including the pagein and pageout primitives described above, the implementation of the fork system call is conceptually very simple. The parent process copies its virtual address space to the child's one page at a time. Note that this may require faulting in the invalid portions of the parent's address space. Since the VAX-11/780 memory-management mechanism can establish only one mapping at a time, the child's address space is actually created indirectly through the kernel virtual memory.
Although conceptually simple, the above scheme has undesirable system performance consequences. Duplication of the parent's private segments generates a sharp and atypical consumption of memory. Since a significant percentage of all forks serve only to create system contexts to be passed to another process via the exec system call, the copying of the parent's private segments is largely unnecessary. The vfork system call, described in §5.1, has been introduced to provide an efficient way to create new system contexts within the current design.
The exec system call, whereby a process overlays its address space also has a simple implementation. The process releases its current virtual memory resources and allocates new ones as determined by the program being executed. Then, the program object file is simply read into the process address space which has been initialized as zero fill on demand pages so as not to generate irrelevant paging from the process' old image.
This implementation suffers from the same problems as the above fork implementation. Initiation of very large programs is very slow, and results in system wide performance degradation due to the loading of the entire program file in the virtual memory before execution commences. A new loader format which allows demand paging from the program object file has been designed to improve large program start up and to eliminate this non-demand situation (see §5.3).
Swapping a process out involves releasing the physical memory currently allocated to it (called the resident set) and writing back its modified pages to its image along with the system maintained state information and page tables. Swapping a process in, on the other hand, involves reading in its page tables and state information and resuming it. Note that as no pages from the process address space are brought in, the process will have to fault them back in as required. The alternative of swapping the resident set in and out is not implemented.
When the amount of available free memory in the system cannot be maintained at a minimal number of free pages by the pageout daemon, then the system invokes the swap scheduler. In order to free memory, the swap scheduler will select a process which is resident and swap it (completely) out. The scheduler prefers first to swap out processes which have been blocked for a significant length of time, and chooses the process which has been in such a state the longest. If there are no such processes, and it is therefore necessary to swap out a process which is or has recently been active, the system chooses from among the remaining processes the one which has been memory resident the longest.
In choosing an active process to swap out, the system checks to guarantee that the process has had a minimal amount of time necessary to demand fault in the number of pages which it had when it was last swapped out (based on maximum expected paging device throughput). This serves to guarantee a minimal amount of progress by a process each time it is swapped in. When a process is forced out while it has many pages, it is given lower priority to return to the set of resident processes than ones which swapped with fewer pages or which are very small.
The swap-in mechanism also recognizes that processes which swapped out with many pages, will need to fault in pages when they are brought in. The system therefore maintains a notion of a global memory deficit, which is the expected short term demand for memory from processes recently brought in, based on the number of pages they were using when they swapped out. The deficit is charged against the free memory available when deciding whether to bring a process in.
In general, this swap scheduling mechanism does not perform well under very heavy load. The system performs much better when memory partitioning can be done by the page replacement algorithm rather than the swap algorithm. If heavy swapping is to occur on moving head devices, then better algorithms could be implemented. High speed specialized paging devices, on the other hand, would suggest different algorithms based on migration.
In a virtual memory environment, handling input/output operations directly to/from process address space without going through the system buffer cache requires special attention. The pages involved in the I/O must be insured to be valid and locked for the duration of the operation. This is accomplished through the virtual segment lock/unlock internal primitives. Locking a virtual segment consists of locking pages that are already valid and faulting/reclaiming invalid pages by simply touching them and refraining from unlocking the page frame (which is allocated in the locked state) after the pagein completion.
In order to allow efficient creation of new processes, a new system primitive vfork has been implemented. After a vfork, the parent and its virtual memory run in the child's system context, which may be manipulated as desired for the new image to be created. When a exec is executed in the child's context, the virtual memory and parent thread of control are returned to the parent's context and a new thread of control and virtual memory are created for the child. This mechanism allows a new context to be created without any copying.
It should be noted that an implementation of fork using a Copy On Write mechanism would be completely transparent and nearly as efficient as vfork. Such a mechanism would rely on more general sharing primitives and data structures than are present in the current version of the system, so it has not been implemented.
In order that efficient random access be permitted in a portable way to large data files, a pair of new system calls has been added: vread and vwrite. These calls resemble the normal UNIX read and write calls, but are potentially much more efficient for sparse and random access to large data files. Vread does not cause all the data which is virtually read to be immediately transferred to the user's address space. Rather, the data can be fetched as required by references, at the system's discretion. At the point of the vread, the system merely computes the disk block numbers of the corresponding pages and stores these in the page tables. Faulting in a page from the file system is thus no more expensive than faulting in a page from the swap device. In both cases all the mapping information is immediately available or can be easily computed from in-core information. Vwrite works with vread to allow efficient updating of large data which is only partially accessed, by rewriting to the file only those pages which have been modified.
Downward compatibility with non-virtual systems is achieved by the fact that read and write calls have the same semantics as vread and vwrite calls; only the efficiency is different. Upward extensibility into a more general sharing scheme is also easy to provide, as vread can be easily simulated by a mapping of the file into the address space with a copy-on-write mechanism on the pages. Although the current mechanism does not share copies of the same page if it is vread twice, the semantics of the system call do not prohibit such an implementation if used with a copy-on-write mechanism. Note that vwrite can also be simulated by a map-out-like mechanism.
The same mechanism that is used to implement the vread system call is used to provide a load format where the pages of the executable image are not pre-loaded, but rather initialized on demand, with the page block numbers only being bound into the page tables at the point of exec. The only change from the other UNIX load formats in this new format is the alignment of data in the load file on a page boundary. Large processes using this format can begin execution very quickly, with page faults causing reading from the executed file.
There are a number of facilities which have not been implemented in the first release of the system as described here.
For example, there are plans to change the system to use 1024 byte disk blocks rather than 512 byte blocks. It has been observed that in many cases the system is limited by the number of disk transactions that can be made per second. Larger disk blocks will help improve disk throughput. On machines with large real memories, using page-pairs in the paging system will also reduce the overhead of the replacement algorithm and increase throughput to the paging device. Since a page contains only 128 words, it does not provide a great deal of locality. It is expected that using 1024 byte pages (in effect) will not degrade the effective memory size significantly.
Another problem associated with the small page size of the VAX architecture is the rate of growth of user page tables.†
.FS † Since a page table entry is 4 bytes, user page tables grow one byte for each 128 bytes of user virtual memory. .FE
For very large processes, this not only results in a significant amount of real memory allocated to page tables, but also increases the system overhead in dealing with them. Effectively supporting extremely large virtual address spaces will require handling translation faults at the first level (i.e., page table faults) whereby real memory for page tables is allocated on demand.
The system changes as presented here are the result of approximately one man year of effort. The base system (a prerelease of UNIX/32V that was maintained as the production system during the paging development) and the paged system consist of approximately 14800 and 16800 total source lines, respectively. The break down of these numbers amongst the various types of source is presented in Table 6.1. Also presented is a comparison that excludes comment lines from the source of the two systems.‡
.FS ‡ For the C source code, the number of occurrences of ``;'' was used as an estimate of the actual number of source lines that were not comments. .FE
We note that the actual number of lines modified to obtain the paged system is considerably more than the simple net difference for each category. In the meantime, for equal configurations, the resident kernel size has increased by about 12K bytes of program and 26K bytes of data resulting in a total size of about 164K bytes (for a 1 megabyte main memory system).
.KF .TS center box; c ||c s ||c s c ||c |c ||c |c c ||c |c ||c |c c ||c |c ||c |c c ||n |n ||n |n. Total Source Lines Non Comment Lines _ _ _ _ Category Base Paged Base Paged System System System System = = = = = Assembly Code 1292 1353 1062 1015 _ _ _ _ _ _ C Code 11581 13405 4891 5614 _ _ _ _ _ Header Files 1997 2068 1223 1316 .TE .ce Table 6.1. Source Code Volume Comparison .sp .KE
The system has been instrumented to collect data related to various paging system activities as well as workload characteristics in general.
Being one of the few quantifiable characteristics of a workload that is also of importance in a virtual memory environment, system-wide distribution of process virtual size was monitored.
.KF .sp 4.15i .ce Fig. 7.1.1. Process size distribution: (a) data, (b) stack segments .sp .25i .KE
The results of integrating process data and stack segment size over user CPU time are shown in Fig. 7.1.1. The two sets of measurements taken almost one month apart indicate an increase from 29.6K to 161.7K bytes and 6.8K to 9.8K bytes for the mean data and stack segment sizes, respectively. The October 18 measurement corresponds to early stages of production run of the system and is representative of the pre-virtual memory workload. The significant increase in the average data segment size within this short period is attributed to the rapid growth of Lisp systems including MACSYMA. The insignificant contribution of the stack segment to total process size is a characteristic of our C intensive workload.
As described in §2, a page can be in three states. Reference to a page in memory but invalid causes a reclaim, whereas reference to one not in memory results in a page-in operation. The service time distributions for these two different types of faults is shown in Fig. 7.2.1.
.KF .sp 4.15i .ce Fig. 7.2.1. Fault service time distribution: (a) reclaim, (b) page-in .sp .25i .KE
For the reclaim time distribution, the first peak corresponds to reclaims from the loop, while the second bump and the long tail are accounted for by the load dependent component of the service time due to reclaiming pages of shared text segments.†
.FS † This operation requires updating the page tables of all processes currently executing the same code, thus varies with load. .FE
Note that the mean reclaim time of 208 microseconds per reclaim represents a negligible delay to user programs. Furthermore, the overall system cost of reclaims through which we simulate the missing reference bits of the architecture has been measured to be less than 0.05% of all CPU cycles.‡
.FS ‡ This cost is actually a function of the paging activity. The number reported here has been averaged over a 28 hour period in a 1.25M byte real memory configuration .FE
The page-in service time distribution is highly load dependent since it includes all of the queueing as well as process rescheduling delays. The configuration with the paging activity on the same arm (an RM-03 equivalent disk) as the temporary and the root file systems results in a 54.9 msec total service time. The significant number of services completed under 20 msec are due to the track buffering capability of the controller being used.
In an effort to compare the performance of the system before and after the addition of virtual memory, a script driven workload was run in a stand-alone manner in both systems under identical configurations consisting of a 1 megabyte main memory, an RP-06 servicing the user file system and an RM-03 shared by the root and temporary file systems in addition to the swapping/paging activity. The swap-based system used for this comparison was quite sophisticated, performing scatter loading of processes into memory and partially swapping processes to obtain free memory.
The basic unit of work generated by the script is made up of four concurrent terminal sessions:
Staggered multiple initiations of from one to four of these terminal sessions were used to create increasing levels of load on the system. Figure 7.3.1 gives the average completion time for each category of session under the two systems.
.KF .sp 7.5c .ce 2 Fig. 7.3.1. Average completion times .br (a) lisp, (b) typeset, (c) ccomp, (d) trivial .sp .KE
For the non-trivial sessions, completion times were very similar under the two systems, with the paging version of the system running (in all but one case) faster, and the swap-based system departing from linear degradation more rapidly. This trend is most noticeable in the response time for the trivial sessions. Systemwide measures collected during the experiments are given in Figure 7.3.2.
.KS .sp 7.5c .ce 2 Fig. 7.3.2. Systemwide measurements .br (a) total (b) average completion time, (c) system time, (d) total page traffic .sp .KE
These measurements show the same trend for both total and average completion times as for individual sessions, with the paging system slightly faster and degrading more linearly than the swap system within the measured range. System overhead was uniformly greater under the paging system, constituting 26 percent of the CPU utilization as compared to 20 percent under the swap system. User CPU utilization was, however, uniformly greater under the paging system, averaging 48 percent, while the swap-based system averaged only 42 percent.
Finally, the total page traffic generated by the two systems was measured. This accounts for both paging and swapping traffic under the paging system, as well as transfer of all system information (control blocks and page tables) under both systems. Although the paging system resulted in far fewer total pages transferred, the actual number of transactions required to accomplish this was much greater since most data transfers, under the paging system, are due to paging rather than swapping activity, and thus take place in very small (512 byte) quantities. We are currently installing modifications in the system to use larger block sizes in both the file and paging subsystems, and expect improved performance from these changes.
Acknowledgments. The cooperation of Bell Laboratories in providing us with an early version of UNIX/32V is greatly appreciated. We would especially like to thank Dr. Charles Roberts of Bell Laboratories for helping us obtain this release, and acknowledge T. B. London and J. F. Reiser for their continuing advice and support.
We are grateful to Domenico Ferrari, Richard Fateman, Jehan-François Pâris, William Rowan, Keith Sklower and Robert Kridle for their participation in the early stages of the design project, and would like to thank our user community for their patience during the system development period.
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