Patent Application: US-201113073407-A

Abstract:
memory demands of large - memory applications continue to remain one step ahead of the improvements in dram capacities of commodity systems . performance of such applications degrades rapidly once the system hits the physical memory limit and starts paging to the local disk . a distributed network - based virtual memory scheme is provided which treats remote memory as another level in the memory hierarchy between very fast local memory and very slow local disks . performance over gigabit ethernet shows significant performance gains over local disk . large memory applications may access potentially unlimited network memory resources without requiring any application or operating system code modifications , relinkling or recompilation . a preferred embodiment employs kernel - level driver software .

Description:
the system according to a preferred embodiment of the present invention is implemented as follows : fig3 shows the client module that handles the paging operations on client - side . that client module has four major components : ( 1 ) the block device interface ( bdi ), ( 2 ) a basic lru - based write - back cache , ( 3 ) a mapping logic that tracks the location of swapped - out pages on different servers , and ( 4 ) a remote memory access protocol ( rmap ) layer , which reliably handles communication between the client and servers . the pager issues read and write requests to the bdi in 4 kb data blocks . the bdi , in turn , performs read and write operations to the lru - based write - back cache . when the cache reaches its capacity , one of the least - recently used dirty pages is evicted to one of the servers using rmap . furthermore , a dirty page is not written back to a server until it is chosen for eviction . fig4 illustrates the server module which handles paging operations on the server side . the server module has two major components : ( 1 ) a hash table that stores client &# 39 ; s pages along with client &# 39 ; s identity ( layer - 2 mac address of the client ), and ( 2 ) the rmap layer , which is similar to the client - side rmap layer . using these two components , the server module blindly stores and retrieves pages for any client in the cluster . once the server reaches capacity , it responds to the requesting client with a negative acknowledgment . it is then the client &# 39 ; s responsibility to select another server to store the page or to write to its local disk , if no other server is available . both client and server modules use hash tables to map memory pages to server and client locations respectively . the hash tables structures themselves consume very little memory . the number of buckets in the hash table remains static upon startup and are allocated using the get free pages ( ) call . linked - lists contained within each bucket hold 64 - byte entry structures that are managed using the linux slab allocator ( which performs fine - grained management of small , same - sized memory objects ). the standard kmalloc ( ) mechanism is not used as it causes unnecessary fragmentation . an important objective of the distributed anemone system is to enable memory - intensive applications to transparently access remote memory resources without requiring any modifications , recompilation or relinking of neither the application nor the kernel . anemone achieves this transparency using a pluggable block device driver , which can provided a block device interface ( bdi ) to the pager as a primary swap partition . page - in / page - out requests can then be sent to the bdi instead of the local disk . because bdi is implemented as a self - contained kernel module , it does not require changing the core os kernel . in fact , one can invoke the open ( ) operation on bdi just like any other standard block device . as a result , the bdi can also be treated as a low - latency store for temporary files and can even be memory - mapped by applications aware of the remote memory . standard disk block devices interact with the kernel through a request queue mechanism , which permits the kernel to group spatially consecutive block i / os ( bio ) together into one “ request ” and schedule them using an elevator algorithm that minimizes disk seek latency . unlike disks , anemone &# 39 ; s read - write latency is independent of the data block being accessed and does not suffer from seek latency overhead . hence anemone &# 39 ; s bdi does not need to group sequential i / o requests together . instead , client module bypasses the request queue and directly intercepts the bios before they are placed on the request queue . hence , the anemone client is able to skip the unnecessary overheads , dispatch the bios in their arrival order , and allow for their completion out of order . bdi maintains its own asynchronous transmission queue to buffer victim pages from the cache . it is best to set this queue size to a value that is just large enough where the pager will not be able to keep it full . however , caution needs to be taken in keeping the memory footprint of this queue at a minimum , usually a value around 1000 pages ( which is a subset of the write - back cache ). rmap is a tailor - made , low - overhead communication protocol which allows remote memory access within the same subnet . rmap implements the following features : ( 1 ) reliable packet delivery , ( 2 ) flow - control , and ( 3 ) fragmentation and reassembly . fig2 illustrates the structure of an rmap packet and its header . while one could technically implement rmap on top of tcp , udp , or even the ip protocol layers , this choice comes burdened with unwanted protocol processing . rmap sits alongside ip but does not use any of ip &# 39 ; s functions . instead rmap takes an integrated , faster approach by communicating directly with the network device driver , sending frames and handling reliability issues in a manner that suites the needs of the anemone system . every rmap message is acknowledged except for keep alive and dynamic discovery messages . timers trigger retransmissions when necessary ( which is rare ) to guarantee reliable delivery . a paging request cannot be allowed to be lost , or the application that depends on that page will fail altogether . rmap also implements flow control to ensure that it does not overwhelm either the receiver or the intermediate network card and switches . however , rmap does not require tcp &# 39 ; s features such as byte - stream abstraction , in - order delivery , or congestion control . hence rmap is implemented as a light - weight window - based reliable datagram protocol . all client nodes keep a static - size window to control the transmission rate , which works very well for a purely in - cluster communication . the last design consideration in rmap is that while the standard memory page size is 4 kb ( or sometimes 8 kb ), the maximum transmission unit ( mtu ) in traditional ethernet networks is limited to 1500 bytes . rmap implements dynamic fragmentation / reassembly for paging traffic . additionally , rmap also has the flexibility to use jumbo frames , which are packets with sizes greater than 1500 bytes ( typically between 8 kb to 16 kb ). jumbo frames enable the rmap to transmit complete 4 kb pages to remote memory using a single packet , without fragmentation . the testbed includes an 8 - port switch that supports jumbo frames ( 9 kb packet size ). a 6 % speed up was observed in rmap throughput by using jumbo frames . all experiments were conducted with 1500 byte mtu with fragmentation / reassembly performed by rmap . the distributed anemone system performs two types of multiplexing in the presence of multiple clients and servers : ( a ) any single client can transparently access memory from multiple servers as one pool via the bdi , and ( b ) any single server can share its unused memory pool among multiple clients simultaneously . this provides the maximum flexibility in efficiently utilizing the global memory pool and avoids resource fragmentation . as servers constantly join or leave the network , the distributed anemone system ( a ) seamlessly absorbs the increase / decrease in cluster - wide memory capacity , insulating the lmas from resource fluctuations and ( b ) allows any server to reclaim part or whole of its contributed memory . this objective is achieved through distributed resource discovery described below . clients can discover newly available remote memory in the cluster and the servers can announce their memory availability . each server periodically broadcasts a resource announcement message ( 1 message every 10 seconds in the prototype ) to advertise its identity and the amount of memory it is willing to contribute to lmas on low - memory clients . each client module maintains a server identity table in which mac addresses of different servers and the memory capacity learned via resource announcements are constantly updated . client module consults this table when selecting a server to perform page - out on a new page . besides resource announcements , servers also piggyback their memory availability information in their page - in / page - out replies to individual clients . this distributed mechanism permits any new server in the network to dynamically announce its presence and allows existing servers to announce their up - to - date memory availability information to clients . distributed anemone also includes soft - state refresh mechanisms to permit clients to track the liveness of servers and vice - versa . firstly , the resource announcement message described above serves an additional purpose of informing the client that the server is alive and accepting paging requests . in the absence of any paging activity , if a client does not receive the server &# 39 ; s resource announcement for three consecutive periods , it assumes that the server is offline and deletes the server &# 39 ; s entries from its server identity table . if the client also had pages stored on that server that went offline , it needs to recover the corresponding pages from a copy stored either on the local disk on another server &# 39 ; s memory . the soft - state refresh mechanism also permits servers to track the liveness of clients whose pages they store . each client periodically transmits a session refresh message to each server that hosts its pages ( 1 message every 10 seconds in the prototype ). each session refresh message carries a client - specific session id . the client module generates a different and unique session id each time the client restarts . if a server does not receive session refresh messages with matching session ids from a client for three consecutive periods , it concludes that the client has failed and frees up any pages stored on that client &# 39 ; s behalf . when a client restarts after failure , its paging requests correspond to a session id that &# 39 ; s different from previous sessions . memory servers themselves are commodity nodes in the network that have their own processing and memory requirements . hence another design goal of the distributed anemone is to a avoid overloading any one server node as far as possible by transparently distributing the paging load evenly . in the earlier centralized architecture , this function was performed by the memory engine which kept track of server utilization levels . distributed anemone implements additional coordination among servers and clients to exchange accurate load information . clients utilize the server load information gathered from resource discovery to decide the server to which they send each new page - out request . this decision process is based upon one of two different criteria : ( 1 ) the number of pages stored at each active server , and ( 2 ) the number of paging requests serviced by each active server . while ( 1 ) attempts to balance the memory usage at each server , ( 2 ) attempts to balance the request processing overhead . there is a small time window between successive resource announcement requests from a server when a client may not have accurate server load information . assume that client sends a new page - out request to a server , and the server finds itself overloaded ( based upon a threshold of memory capacity or request throughput ). in this case , the server sends a negative acknowledgment to the client , declining to store the new page . the client can then either select a different server or store the page to local disk , if no other server is available . the ultimate consequence of failure in swapping to remote memory is no worse than failure in swapping to local disk . however , the probability of failure is greater in a lan environment because of multiple components involved in the process , such as network cards , connectors , switches etc . at the communication protocol level , anemone implements the rmap protocol to ensure reliable transfer of memory pages across the network . additionally , two alternatives for tolerating server failures are proposed : ( 1 ) to maintain a local disk - based copy of every memory page swapped out over the network . this provides same level of reliability as disk - based paging , but risks performance interference from local disk activity ; and ( 2 ) to keep redundant copies of each page on multiple remote servers . this approach avoids disk activity and reduces recovery - time , but consumes bandwidth , reduces global memory pool and is susceptible to network failures . the distributed anemone prototype was evaluated , with the goal answering the following key questions : ( 1 ) latency : what reduction in paging latency does anemone deliver when compared to disk - based paging ? how do the latency distributions of anemone vary across sequential / random access patterns ? ( 2 ) speedup : what application speedups can be obtained with anemone for real - world unmodified applications ? how do the speedups compare against using local memory and against swapping to disk ? how do the speedups vary when executing multiple lmas within a single client ? ( 4 ) protocol performance : to what extent do the overheads due to rmap affect paging performance ? the results can be summarized as follows . distributed anemone reduces read latencies to an average 160 μs compared to 8 . 3 ms average for disk and 500 μs average for centralized anemone . for writes , both disk and anemone deliver similar latencies due to write caching . in these experiments , anemone delivers a factor of 1 . 5 to 4 speedup for single process applications , and delivers up to a factor of 14 speedup for multiple concurrent applications . the distributed anemone system can successfully operate in the presence of both multiple clients and multiple servers . the experimental testbed consisted of one 64 - bit low - memory amd 2 . 0 ghz client machine containing 256 mb of main memory and nine remote - memory servers . the servers consisted of : four 512 mb machines , three 1 gb machines , one 2 gb machine , one 3 gb machine . the 512 mb servers range from 1 . 7 ghz to 800 mhz intel processors . the other 5 machines are all 2 . 7 ghz and above intel xeons . one of the 1 gb memory machines is also a 64 - bit machine , which is used to perform experiments with multiple - client workloads on the system as described later . when active , the cluster contains a total of about 9 . 6 gb of usable memory ( after accounting for local memory usage within each machine ). for disk based tests , a western digital wd800jd 80 gb sata disk , with 7200 rpm speed , 8 mb of cache and 8 . 9 ms average seek time was used , ( which is consistent with the results ). this disk has a 10 gb swap partition reserved on it to match the equivalent amount of remote memory available in the cluster , which is used exclusively when comparing the distributed anemone system against the disk . each machine is equipped with an intel pro / 1000mt gigabit ethernet card connected to two 8 - port gigabit switches , one from netgear and one from smc . a page - in operation is more time - critical than page - out operation because the application cannot proceed without the required page . the various components of the latency in processing a 4 kb page - in request from the pager involves the following steps : ( 1 ) the client module looks up the offset - to - server mapping in its mapping table . ( 2 ) the client transmits a request for the page directly to the server . ( 3 ) the server looks up the page for that offset and transmits the requested page directly back to the client . ( 4 ) the client module returns the page back to the pager and the process resumes . the entire sequence requires only 2 network transmissions and only a few microseconds of processing time . again , on a per - latency basis , a single page - in request takes 160 microseconds average with a standard deviation of 22 μs , which is about 50 times faster than page - in requests to disk . the distribution of observed read and write latencies was next examined for sequential and random access patterns . fig5 compares the cumulative distributions of latencies with disk and anemone for random and sequential read requests . similarly , fig6 compares the two for random and sequential write requests . though real - world applications rarely generate purely sequential or completely random memory access patterns , these graphs provide a useful measure to understand the underlying factors that impact application execution times . for random read requests in fig5 , most requests to disk experience a latency between 5 to 10 milliseconds . on the other hand most requests in anemone experience only around 160 microsecond latency . for sequential read requests in fig5 , disk shows a slightly superior latency distribution than anemone . most sequential requests are serviced by the on - board disk cache within 3 to 5 microseconds because sequential read accesses fit well with the motion of disk head , eliminating seek and rotational overheads . in contrast , anemone still delivers a range of latency values , most below 100 microseconds . this is because network communication latency dominates in anemone even for sequential requests , though it is masked to some extent by the prefetching performed by the pager and file - system . the write latency distributions for both disk and anemone in fig6 are comparable to each other , with most latencies being close to 9 microseconds because writes typically return after writing to the disk cache or anemone &# 39 ; s client cache . the first application tested is a graphics rendering program called pov - ray [ 20 ]. the memory consumption of pov - ray was varied by rendering different scenes with increasing number of spheres . fig7 shows the completion times of these increasingly large renderings up to 3 . 4 gb of memory versus the disk using an equal amount of local swap space . fig7 clearly shows that anemone delivers increasing application speedups with increasing memory usage and is able to improve the execution time of a single - process pov - ray by a factor of 4 for 3 . 4 gb memory usage . the second application is a large in - memory quicksort program that uses an stl based implementation from sgi [ 21 ], with a complexity of o ( n logn ) comparisons . randomly populated large in - memory arrays of integers were sorted . fig8 shows that anemone delivers speedup of factor of 2 . 4 for single - process quicksort having 5 gb memory usage . the third application is the popular network simulator ns2 [ 26 ]. a delay partitioning algorithm on a 6 - hop wide - area network path using voice - over - ip traffic traces was simulated . with ns2 requiring 1 gb memory , anemone speeds up the simulation by a factor of 4 compared to disk based paging . the fourth application is the k - nearest neighbor ( knn ) search algorithm on large 3d datasets , which are useful in applications such as medical imaging , molecular biology , cad / cam , and multimedia databases . when executing knn search algorithm over a dataset of 2 million points consuming 1 . 5 gb memory , distributed anemone speeds up the simulation by a factor of 3 . 7 over disk based paging . the performance of distributed anemone system was tested under varying levels of concurrent application execution . multiple concurrently executing memory - intensive processes tend to stress the system by competing for computation , memory and i / o resources and by disrupting any sequentiality in the paging activity . fig9 and 10 show the execution time comparison of anemone and disk as the number of pov - ray and quicksort processes increases . the execution time measures the time interval between the start of the multiple process execution and the completion of last process in the set . each process consumes 100 mb of memory . fig9 and 10 show that the execution times using disk - based swap increases steeply with number of concurrent processes . the paging activity loses sequentiality with increasing number of processes , making the disk seek and rotational overheads dominant . on the other hand , anemone reacts very well to concurrent system activity and the total execution time increases at a much slower pace . this is because , unlike disk based paging , anemone encounters a steady paging latency over the network even as the paging activity loses sequentiality of access . with 12 - 18 concurrent memory - intensive processes , anemone achieves speedups of a factor of 14 for pov - ray and a factor of 6 . 0 for quicksort . one of the important knobs in rmap &# 39 ; s flow control mechanism is the client &# 39 ; s transmission window size . using a 1 gb quicksort , fig1 shows the effect of changing this window size on three characteristics of the anemone &# 39 ; s performance : ( 1 ) the number of retransmissions , ( 2 ) paging bandwidth , which is represented in terms of “ goodput ”, i . e . the amount of bandwidth obtained after excluding retransmitted bytes and ( 3 ) application completion time . as the window size increases , the number of retransmissions increases because the number of packets that can potentially be delivered back - to - back also increases . for larger window sizes , the paging bandwidth is also seen to increase and saturates because the transmission link remains busy more often delivering higher “ goodput ”, in spite of an initial increase in number of retransmissions . however , if driven too high , the window size will cause the paging bandwidth to decline considerably due to increasing number packet drops and retransmissions . the application completion times depend upon the paging bandwidth . initially , an increase in window size increases the paging bandwidth and lowers the completion times . similarly , if driven too high , the window size causes more packet drops , more retransmissions , lower paging bandwidth and higher completion times . to measure the control traffic overhead due to rmap , the percentage of control bytes generated by rmap compared to the amount of data bytes transferred while executing a 1 gb pov - ray application is measured . control traffic refers to the page headers , acknowledgments , resource announcement messages , and soft - state refresh messages . the number of servers is first varied from 1 to 6 , while a single client was executing the pov - ray application . next , the number of clients is varied from 1 to 4 ( each executing one instance of pov - ray ), with 3 memory servers . the percentage overhead of control traffic was consistently measured to be around 1 . 74 %— a very small overhead for the amount of page data transmitted . although this invention has been illustrated by reference to specific embodiments , it will be apparent to those skilled in the art that various changes and modifications may be made which clearly fall within the scope of the invention . the invention is intended to be protected broadly within the spirit and scope of the appended claims . a . acharya and s . setia . availability and utility of idle memory in workstation clusters . in measurement and modeling of computer systems , pages 35 - 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