text:
|Publication number
|Publication type
|Application number
|Publication date
|May 17, 2011
|Filing date
|May 8, 2006
|Priority date
|May 8, 2006
|Fee status
|Also published as
|Publication number
|Sorin Faibish, Stephen A. Fridella, Xiaoye Jiang, Uday K. Gupta
|Original Assignee
|Emc Corporation
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|BiBTeX, EndNote, RefMan
|Patent Citations (30), Non-Patent Citations (16), Referenced by (19), Classifications (11), Legal Events (2)
|External Links: USPTO, USPTO Assignment, Espacenet
The present invention generally relates to computer data storage systems, and more particularly to network file servers. The present invention specifically relates to a data processing system in which a first processor distributes data blocks to a second processor for use in creating or extending files of a file system.
Mainframe data processing, and more recently distributed computing, have required increasingly large amounts of data storage. This data storage has been most economically provided by an array of low-cost disk drives integrated with a large semiconductor cache memory.
Many applications access data that is stored in files of a file system. The file system is a hierarchical arrangement of the files such that a particular file has a path name through the file system. The path name includes a sequence of names beginning with the name of a root directory of the file, possibly including the names of one or more sub-directories, and ending with a name for the file. The file system also includes a mapping of the directories and files in the file system to file system blocks, and a mapping of the file system blocks to logical blocks of storage.
In order to access data of a file, an application program sends one or more file access requests to a system program known as a file system manager. The application program typically specifies a pathname for the file, an offset from the beginning of the file, and an extent of data to be read or written starting at the offset in the file. For example, the offset is a specified number of bytes from the beginning of the file, and the extent of data is a specified number of consecutive bytes starting at the offset in the file. The file system manager maps the file name, offset, and extent to logical file system blocks mapped to logical blocks of storage, and sends logical block access commands to a volume manager or storage subsystem.
The volume manager or storage subsystem maps the logical blocks of storage to physical blocks of storage in storage devices such as disk drives. This logical-to-physical mapping may involve striping of the extent of data across a redundancy group of disk drives in a Redundant Array of Inexpensive Disks (RAID).
In the data processing art, there has been a general trend to specialize and distribute data processing tasks among an increasing number data processors. For example, more than a decade ago, it was common for a mainframe computer to be programmed with a file system manager and various application programs that invoked the file system manger in order to access files of the file systems. In turn, the mainframe computer sent logical block access commands to another mainframe processor of a cached disk array.
More recently, it has been common for application programs to be executed by workstations such as personal computers networked to file servers. Each file server is programmed with a file system manager. Each file server may include a volume manager for access to storage of disk drives in the file server. However, file servers have been networked or clustered in various ways to enable share access to storage subsystems or arrays of disk drives by multiple workstations.
Data consistency problems may arise if two file servers share access to the same file system in storage. As described in Xu et al. U.S. Pat. No. 6,324,581, one way to solve this data consistency problem is to designate one of the file servers to be an exclusive owner of access rights to each file system. The exclusive owner of the access rights to a file system, however, may delegate data access or metadata management tasks to other file servers. For example, if a first file server receives a request from a network client for access to a file system owned by a second file server, then the first file server sends a metadata request to the second file server. The second file server responds by placing a lock on the file and returning metadata of the file. The first file server uses the metadata of the file to formulate a data access command that is used to access the file data in the file system directly to the disk array over a bypass data path that bypasses the second file server.
As further described in Jiang et al. U.S. patent application Publication 2005/0240628 published Oct. 27, 2005, metadata management in a file server or storage network is delegated from a primary data processor to a secondary data processor in order to reduce data traffic between the primary data processor and the secondary data processor. The primary data processor retains responsibility for managing locks upon objects in the file system that it owns, and also retains responsibility for allocation of free blocks and inodes of the file system. The leasing of free blocks and inodes to the secondary and the granting of locks to the secondary enables the secondary to perform other metadata management tasks such as appending blocks to a file, truncating a file, creating a file, and deleting a file.
In accordance with one aspect, the present invention provides a method of operating a primary data processor and a secondary data processor for access to a file system in data storage. The method includes the primary data processor pre-allocating data blocks of the data storage to the file system before the pre-allocated data blocks are needed for extending the file system, and the secondary data processor using at least some of the pre-allocated data blocks for extending the file system. The method further includes selecting an address mapping for at least some of the pre-allocated blocks in order to reduce block scatter on disk and improve I/O performance.
In accordance with another aspect, the invention provides a method of operating a primary data processor and a secondary data processor for access to a file system in data storage. The method includes the primary data processor pre-allocating data blocks of the data storage to the file system before the pre-allocated data blocks are needed by the secondary data processor for extending the file system, and the secondary data processor using at least some of the pre-allocated data blocks for extending the file system. The primary data processor pre-allocates an initial chunk size of data blocks when the file system is created, and upon expiration of a release timeout duration from the pre-allocation of the initial chunk size, the secondary data processor releases to the primary data processor an excess of the pre-allocated data blocks over a release threshold.
In accordance with still another aspect of the present invention, a data processing system includes data storage, a primary data processor programmed for pre-allocating data blocks of the data storage to a file system in the data storage before the pre-allocated data blocks are needed for extending the file system, and a secondary data processor coupled to the data storage for access to the file system in the data storage and coupled to the primary data processor for obtaining a map of the data blocks pre-allocated to the file system. The data processing system is further programmed for selecting an address mapping of at least some of the pre-allocated data blocks in order to reduce block scatter on disk and improve I/O performance.
Additional features and advantages of the invention will be described below with reference to the drawings, in which:
While the invention is susceptible to various modifications and alternative forms, a specific embodiment thereof has been shown in the drawings and will be described in detail. It should be understood, however, that it is not intended to limit the invention to the particular form shown, but on the contrary, the intention is to cover all modifications, equivalents, and alternatives falling within the scope of the invention as defined by the appended claims.
With reference to
The network file server 24 includes a cached disk array 28 and a number of data mover computers, for example 25, 26, 27, and more. The network file server 24 is managed as a dedicated network appliance, integrated with popular network file systems in a way, which, other than its superior performance, is transparent to the end user. The clustering of the data movers 25, 26, 27 as a front end to the cache disk array 28 provides parallelism and scalability. Each of the data movers 25, 26, 27 is a high-end commodity computer, providing the highest performance appropriate for a data mover at the lowest cost. The network file server 24 also has a control station 29 enabling a system administrator 30 to configure and control the file server.
In the network environment of
The cache disk array 28 includes a cache memory 34 for storing most recently accessed logical data blocks. If a logical data block to be accessed is not found in the cache memory, then the cache disk array performs a logical-to-physical mapping 35 to recall the data block from disk storage of the file systems 31, 32, 33. The logical-to-physical mapping 35, for example, may stripe contiguous data of a file across a redundancy group of disk drives in accordance with a particular RAID level.
If a data mover receives a client request for read-write access to a file in a file system that the data mover does not own, then the data mover owning the file system is said to be primary with respect to access to the file, and the data mover having initially received the request from the client is said to be secondary with respect to access to the file. In this situation, the secondary data mover can either forward the request to the primary data mover, or the secondary data mover can request a lock upon the file from the primary data mover, and access the file once the primary data mover grants the lock to the secondary data mover. The forwarding technique is preferred when the read-write access involves a relatively small amount of data to be read or written. Otherwise, it is preferred for the secondary data mover to access the file directly once the primary data mover grants a lock to the secondary data mover.
In the network file server 24, the secondary data mover can also request metadata for access to a file owned by the primary data mover. In response to the metadata request, upon granting the lock upon the file, the primary data mover returns to the secondary data mover metadata specifying data storage locations in the cached disk array 28 for storing data of the file. The secondary data mover uses the metadata to produce at least one data access command for accessing the data storage locations in the cached disk array for storing data of the file. For a write operation, the secondary data mover may modify the metadata. When the secondary data mover is finished writing to the file, the secondary data mover returns any modified metadata to the primary data mover. See, for example, Xu et al. U.S. Pat. No. 6,324,581.
As described in Jiang et al. U.S. Patent Publication 2005/0240628, the primary data mover can also lease free file system inodes and free data blocks to the secondary data mover so that the secondary data mover can perform metadata management tasks such as appending blocks to a file, creating a file, modifying a file, and deleting a file.
The data mover 25 has a network interface 200 for communication of data packets over the data network (21 in
If the UxFS module 205 receives a request to access a file system owned by another data mover, then it access a Multiplex File System (MPFS) module 206. The MPFS module 206 decides whether to forward the file access request to the data mover owner of the file system, or to request metadata of the file from the data mover owner of the file system, so that the UxFS module 205 may use the metadata for accessing the file.
The UxFS module 205 accesses data organized into logical volumes defined by a module 207. Each logical volume maps to contiguous logical storage addresses in the cached disk array 28. The module 207 is layered over a SCSI driver 208 and a Fibre-Channel protocol (FCP) driver 209. The data mover 25 sends storage access requests through a host bus adapter 210 using the SCSI protocol, the iSCSI protocol, or the Fibre-Channel protocol, depending on the physical link between the data mover 25 and the cached disk array 28.
The primary data mover 26 includes modules 220 to 230 that are similar to respective modules 200 to 210 of the secondary data mover 25.
In accordance with an aspect of the present invention, the secondary data mover 25 is programmed with a File Mapping Protocol (FMP) client module 243 for obtaining an allocation of free data blocks from the primary data mover. The primary data mover has a FMP server 247 that creates and manages a file system used as a pool of free data blocks that have been reserved for allocation to file systems that are owned by the primary data mover. The FMP server 247 responds to a request for data blocks from the FMP client by allocating free pool file system data blocks to the secondary data mover and returning a map of these allocated pool file system data blocks to the FMP client. The map of these allocated data blocks, for example, is in the form of list of logical offset/range in logical volumes on the cached disk array 28.
The FMP server 247 maintains a block allocation map 248 indicating the data blocks that have been reserved for use by the primary data mover and which of these data blocks have been leased out to each secondary data mover. This pool file system includes a multiplicity of sparse files, which the primary data mover may allocate to respective secondary servers or external hosts to use as a source of file system inodes and file system data blocks or as logical raw volumes on which to build UxFS or other host file systems.
The FMP client 243 of the secondary data mover 25 maintains a pool file system map 244 indicating the files and data blocks that have been allocated by the primary data mover to the secondary data mover. The FMP client 243 also has the ability to select data blocks from this pool so that the selected data blocks will have physical addresses especially suited to the type of access to be performed by the secondary data mover 25. Because the selected data blocks are to be linked into particular files, the selected data blocks should also have particular logical addresses. Thus, the FMP client also is provided with an ability to select and set a desired logical-to-physical address mapping for the selected data blocks.
When allocating data blocks to a file or file system, the FMP client 243 can adjust the logical-to-physical mapping of the selected data blocks independent of the data block mapping by the UxFS module 205 and independent of the logical-to-physical mapping of the logical volumes provided by the logical volumes module 207. For this purpose, the secondary data mover 25 is programmed with a second CFS module 241 and a second VFS module 242 layered below the UxFS module 205 and above the logical volumes module 207. In effect, the CFS module 241 and VFS module 242 provide a logical-to-logical address translation or re-mapping for the logical data blocks.
Typically the FMP server 247 will respond to a request from the FMP client for free data blocks by filling the pool FS map with many more data blocks than are presently needed. After a certain timeout a garbage collector daemon 245 checks whether this excess capacity has dropped to a certain release threshold, and if not, the excess allocated capacity over the release threshold is released back to the pool FS map, and the released data blocks are given back to the primary data mover owner of the pool file system. In this fashion, the garbage collection process ensures that the pool file system is not wastefully utilized.
A thin provisioning module 246 is provided to support sparse files. The thin provisioning module 246 keeps track of data blocks that are not yet actually used in sparse files. The thin provisioning module 246 may adjust a reserve threshold of data blocks that should be kept in the pool file system to support sparse files. The FMP client 243 requests the primary data mover to replenish the pool file system when the number of data blocks in the pool file system falls below the reserve threshold. The thin provisioning module 246 may pre-allocate data blocks of the pool FS map to the data blocks that are not yet actually used in the sparse files when data blocks are available, and may release such pre-allocated data blocks when there is a need for free data blocks that is not met by the primary data mover when the FMP client requests the primary data mover to replenish the pool file system when the number of data blocks in the pool file system falls below the reserve threshold. The primary server 26 can also revoke or request back pre allocated blocks unused by the secondary server to maintain the reserve threshold at a given level, before asking for more disk capacity to the storage array.
The storage allocation server 253 divides the free storage up into disjoint sets and allocates a respective one of the disjoint sets to each of the metadata servers for use in extending the respective file systems owned by each of the metadata servers. Thus, the storage allocation server 253 delivers a respective map of the free storage to each metadata server 250, 251, 252 without having to worry about data consistency because each metadata server is given access to a respective disjoint set of disk blocks so locking conflicts do not arise between the metadata servers.
The MPFS client/servers 247, 248, the metadata servers 250, 251, 252, and the storage allocation server 253 are linked via the storage area network 254 to a number of storage arrays 255 containing disk drives. By using the iSCSI protocol, the storage area network 254 can be extended to a wide area network making this cluster architecture scalable without limits.
The MPFS client/servers 247, 248 read and write client data over the storage area network 254 to the disk drives in the storage arrays 255, 256 at physical addresses (e.g., at disk blocks of 512 bytes) specified by metadata obtained from the metadata servers 250, 251, 252 acting as storage allocation servers. The metadata servers commit metadata changes to the on-disk file systems by transmitting the metadata changes over the storage area network 254 to the disk drives in the storage arrays 255, 256. For recovery purposes, the storage allocation server 253 keeps an on-disk record of the respective disjoint sets of disk blocks allocated to each of the metadata servers.
In the system of
The storage allocation server 253 maintains a local map 308 of free storage and storage allocated to each metadata server, and allocates this free storage to the metadata servers as needed. It may as well serve metadata to other MPFS clients. The local map 308 need not identify all of the free storage available via the storage area network 254. Also, once a metadata server commits its allocated storage to a file system that it owns, the storage allocation server need not continue to keep a record of the allocated storage for recovery purposes. Thus, the metadata server returns a commit message indicating that the storage allocated to it is no longer free, and the storage allocation server responds to this commit message by removing the committed storage from its free storage blocks in map 308.
The allocation of free storage blocks by a primary server to a secondary server in the system of
When the storage allocation server determines that a metadata server has failed, any uncommitted blocks processed by the metadata server are lost. Therefore, the allocation server can access its local block allocation map to reclaim the not yet used blocks that were allocated to the failed metadata server. If the storage allocation server fails, the metadata servers can continue processing with the not yet used blocks that have been allocated to them.
The storage allocation server should keep its local map 308 full of sufficient free storage for satisfying promptly requests for allocation from the metadata servers, and the storage allocation server should allocate sufficiently large chunks of storage to the metadata servers so that as data blocks are allocated to files, the data blocks will not be scattered poorly on the physical disk. For example, the storage allocation server allocates large chunks of storage to a respective sparse file for each metadata server. This sparse file serves as a storage pool or logical volume container for storage allocated to the metadata server. Thus, as shown in
By maintaining each pool file system file allocated to the metadata server 250, 251 or 252 as a sparse file, the metadata server need not have the entire map of the pool file system blocks all of the time, and the number of blocks allocated to the pool file system may increase as blocks are written to the file systems owned by the metadata servers. The structure of the file systems owned by the metadata servers, however, should be aligned to the structure of the pool file system. Otherwise, the blocks of the file systems owned by the metadata server will possibly be fragmented on the pool file system and as a result scattered very badly on the physical disk resulting in poor I/O performance.
In order to solve this alignment problem, the FMP client 303 requests a large number of pool FS blocks to be allocated; for example, for each new UxFS file system file opened by the metadata server 250, an entire Cylinder Group of the pool file system will be requested from the storage allocation server, even if it may only use one single UxFS file system block at the time of the initial request. Because the pool file system is also a UxFS file system, the secondary file system will grow using the same structure as the pool file system. After a certain timeout the garbage collector daemon 305 releases the unused pool file system blocks allocated to the secondary server and not used. This garbage collection process ensures that the pool file system is properly utilized not wastefully. In a similar manner, when a sparse file is used as a raw volume, the storage allocation server allocates file system blocks in a structure matching the structure of the file system laid on the raw logical volume.
The FMP client 303 adapts to the structure of the raw volume layout and allocates blocks in the pool file system according to the behavior of the file systems that it owns. If the metadata server supports different kinds of file systems (such as ext3 or xfs in addition to UxFS), then the behavior depends on the particular kind of file system. In this case, it may be desirable for the pool file system to include, for each metadata server, a separate file of pre-allocated blocks for each different kind of file system. Therefore it is possible to improve performance of file systems that do not possess good block allocation techniques by improving the layout of the data blocks on the physical disk for enhanced I/O performance.
It is also possible that different file systems of the same type but used for different applications or system tasks (such as an electronic mail archiving application, or a seismic data processing application, or a database application, or a snapshot copy task) may exhibit different behavior. In this case, it also may be desirable for the pool file system to include, for each metadata server, a separate file of pre-allocated blocks for each file system for the different kind of application or system task. An adaptive mechanism can be used to learn the behavior of the particular kind of file system or application and improve the layout on disk by matching the pre-allocation to the structure of the file system or application I/O pattern.
The file system access at the upper level 341 causes a file system manager at a next lower level 342 to select, map and link inodes and file system data blocks in accordance with the kind of file system being accessed. For example, for a UNIX-based file system, the inodes and file system data blocks are 8 K bytes in size. The file system manager accesses logical blocks, which are normally mapped to physical blocks at a lower level 345. The mapping at the lower level 345 is typically performed for a certain level of RAID redundancy or is performed in a storage subsystem such as a cached disk array. The mapping at the lower level 345 results in physical addresses for addressing disk storage at the bottom level 356. Typically each physical address specifies a SCSI port and a track, sector, and head for a disk block containing 512 bytes.
Selective mapping of incrementally added data blocks occurs at a level 343 between the level 342 of the file system manager and the level 345 of the logical-to-physical mapping. The selective mapping at the level 343 is performed to improve layout of the file system on disk and avoid block scatter so as to improve I/O performance by reducing disk drive head excursions for typical access patterns of the applications and system tasks. Logical blocks that are accessed in a contiguous fashion when reading or writing to files should be mapped to contiguous physical blocks on disk.
The selection of the mapping to use at the level 343 is essentially a de-fragmentation process. The selected mapping is then programmed into the combination of the second CFS layer 241 and the second VFS layer 242 in
The selective mapping at the level 343 is performed in conjunction with a process 344 of adaptive pre-allocation of data blocks for avoiding delay in allocation and in the selection of the mapping to be used at level 343. For example, a respective sparse file in the pool file system and respective allocation parameters (such as the reserve threshold, release threshold, release timeout, and pre-allocation chunk size) are provided for different kinds of file systems and applications and system tasks.
For example, to begin building a new file system, the allocation server allocates an entire cylinder group of data blocks to the metadata server. The size of the cylinder group depends on the particular kind of file system. For example, a UNIX cylinder group contains 8,192 file system blocks, and each file system block contains 8 K bytes. When the metadata server requests additional blocks for this UNIX-based file system, the allocation server allocates an additional half cylinder group to the metadata server. Thus, as shown in
The FMP database of
The FMP database of
If in step 351 a new file system is not needed and after step 353, execution continues to step 356. In step 356, if a file is not to be created or extended, then execution loops back to step 351. Otherwise, execution continues to step 357. In step 357, if the sparse file for the file system has a sufficient number of not yet used blocks for creating or extending the file, then execution continues to step 358 to get these needed blocks from the sparse file, link them into the file system or file, and commit them to the on-disk file system. (This involves updating the pool file system map so that blocks are no longer indicated as "not yet used", for example by updating the pool file system map to remove the blocks from a list of offset/range of not yet used blocks.) In step 359, if the number of not yet used blocks in the sparse file is not less than the reserve threshold, then execution loops back to step 351. Otherwise, execution continues to step 360 to get a pre-allocation chunk of blocks from the storage allocation server and to reset the release timer. After step 360, execution loops back to step 351.
In step 357, if there is not a sufficient number of not yet used blocks in the sparse file, then execution branches to step 361. In step 361, the needed blocks are obtained from the storage allocation server, plus additional blocks, the reserve and release thresholds are increased, and the release timer is reset. For example, the reserve threshold and the release threshold are increased by a fraction of the number of blocks that are needed less the number of not yet used blocks in the sparse file, and the number of blocks obtained in step 361 includes a sufficient number of pre-allocation chunks to exceed the number of blocks that are needed less the number of not yet used blocks in the sparse file plus the updated release threshold. After step 361, execution loop back to step 351.
The selective decreasing of the allocation chunk size in step 355 and increasing of the reserve and release thresholds in step 361 provides fast adaptation to an over supply and under supply of not yet used blocks. The allocation chunk size and the reserve and release thresholds should be brought back slowly to their initial default values. These initial default values themselves can be adjusted more slowly based on the allocation history, for example, in order to maintain a desired relationship of the pre-allocation chunk size, release threshold, and reserve threshold with respect to statistics of the number of not yet used blocks allocated to the metadata server, as generally shown in
Typically when a file is created in an UxFS file system, there is a default initial file allocation. For example, for a file created in response to a NFS request from a network client, the initial default size is 32 file system blocks of 8 K bytes per block. It is preferred for the metadata server to initially allocate a chunk of blocks that is a multiple of the initial default size, such as 64 times the initial default size, although the multiple could be selected based on the particular application that is requesting the file to be created (for example, as indicated by a file name extension identifying a file type). Blocks pre-allocated to the file in excess of a release threshold could be returned to the sparse file of the pool file system after a release timeout interval following the creation of the file. The multiple blocks should be contiguous on disk to the extent reasonably possible, and the additional blocks in each chunk should be reserved for extending the file. Thus, there should be a pre-allocation of physically contiguous disk blocks to files at least for applications that are expected to extend the files after the files are created and initially written to.
For file systems and applications permitting multiple writers (e.g., application processes or system processes) to write to a sparse file, it may also be desirable to retain a history of where each writer has been writing to the file and how much data has been allocated to each writer in order to pre-allocate physically contiguous file system data blocks for each writer to continue writing in a logically contiguous fashion.
To support a sparse file as shown in
A storage allocation server can also be programmed to function as a primary server for managing metadata of a production dataset and snapshot copies of the production dataset while delegating to a secondary storage server the read-write access to the production dataset. This can be done in such a way that most of the management of the snapshot metadata can be done "out of band" with respect to the read-write access to the production dataset. Therefore the I/O performance of the secondary storage server is not reduced very much by the snapshot copy process.
The block transitions to a next state S4 when a snapshot is taken. In state S4, the block is shared between the production file and the most recent snapshot copy of the production file. The block transitions to a next state S5 when the storage server performs a first write to the block since the most recent snapshot of the production file. In state S5, the storage block is in the snapshot copy only and is no longer in the production file. The new data for the file block is not stored in this storage block and instead is stored in another storage block that has been pre-allocated to the file. The storage block in state S5 remains in state S5 until the storage allocation server moves or discards the snapshot copy data from the storage block. Because the storage block is co-located with other storage blocks of the file, it often is desirable for the storage allocation server to move or discard snapshot data in order to pre-allocate the storage block for re-use in the file. In this case, the block transitions from state S5 back to state S0, and then once again to state S1 after the appropriate allocation operation.
For example, to create a "write anywhere" snapshot of a file, the file's metadata is made "read-only." Then the inode of the file is cloned to create a production file inode and a snapshot file inode. Initially, the indirect block tree of the file is linked to both of these inodes. When new data is first written to a block of the production file since the time of the snapshot, the new data is written to a newly allocated block, and the block pointer to the original data block is changed in the production file inode so that it points to the newly allocated block, and one bit in this block pointer indicates that this block has been written to since the time of the snapshot. For keeping a chronological series of snapshots of the file, this one bit is more generally used as an ownership bit indicating whether or not the data of the pointed-to data block changed prior to the time of the snapshot and after the time of the next oldest snapshot. Further details regarding this procedure of creating and maintaining write-anywhere snapshots of a file are found in Bixby et al., U.S. patent application Pub. No. 2005/0065986 published Mar. 24, 2005 entitled "Maintenance of a File Version Set Including Read-Only and Read-Write Snapshot Copies of a Production File," incorporated herein by reference.
The secondary storage server 402 receives read/write requests and data from an application program 407 executed by the client 401. The production file itself can be a container for a UxFS file system that is accessed by the application program 407. In this case, the logical extent of the production file serves as a logical volume upon which the UxFS file system is built. By using a single production file as a container for a UxFS file system, it is possible for a conventional backup, migration, or replication program to perform a single backup, migration, or replication upon an entire UxFS file system including multiple files. The container file serves as a convenient package for the multiple files, and the attributes of the container file will indicate when the version of the UxFS file system in the container file was created and last accessed. Further details regarding the use of a file as a logical volume for a UxFS file system are found in Virendra M. Mane, "File Based Volumes and File Systems," U.S. patent application Ser. No. 11/301,975 filed Dec. 13, 2005, incorporated herein by reference.
The secondary storage server 402 has a local cache 408 of production file mapping information. This production file mapping information includes respective storage addresses for file blocks, and also block state information indicating whether or not each production file block in local cache 408 is either in a "READ-ONLY" state or a "READ-WRITE" state. A respective flag can be used to indicate the state of each file block in the local cache, or the local cache can be subdivided into a list of "READ-ONLY" blocks and a list of "READ-WRITE" blocks. For example, when a snapshot is taken, the state of each of the production file blocks in the local cache is set to "READ-ONLY" because each block is then shared between the production file and this most recent snapshot until a write to the block occurs. The secondary storage server 402 also has a write map 409 of blocks that have been pre-allocated to the production file.
At any particular time, the local cache 408 and the write map 409 need not include mapping information for all of the blocks in the file. Typically the local cache 408 and write map 409 are initially empty. From an initially empty state, the local cache 408 of production file mapping information becomes filled in response to read or write requests from the application program 407. When production file mapping information is needed for responding to a read request, the secondary storage server 402 sends a "Get Map" request to the storage allocation server 403, and the storage allocation server returns the requested production file mapping information.
From an initially empty state, mapping information about pre-allocated blocks is added to the write map 409 in response to write requests from the application program 407 for the first writes to file blocks since the last snapshot. When mapping information about a pre-allocated block is needed for responding to a write request, the secondary storage server 402 sends an "Allocate Space" request to the storage allocation server 403, and the storage allocation server returns the requested mapping information. The secondary storage server 402, however, may request or the storage allocation server 403 may provide mapping information about more blocks than needed for one write operation. For example, blocks may be pre-allocated in response to an explicit request from the application program to fill holes in the file or to extend the file.
In a similar fashion, the storage allocation server 403 has a local cache 410 of production file and snapshot mapping information, and a local map 411 of pre-allocated blocks. The local cache 410 of production file and snapshot mapping information is used to maintain an in-cache file system including the production file and its snapshots, corresponding to an on-disk file system of file metadata 406 in the storage 404. The local map 411 of pre-allocated blocks permits the storage allocation server to pre-allocate selected storage blocks to particular offsets in the production file well in advance of receipt of "Allocate Space" requests for the selected storage blocks. In this case, the local map 411 may include an indication of whether or not the metadata information for the pre-allocated blocks has been sent to a secondary storage server, and if so, an identification of the secondary storage server.
The secondary storage server 402 sends read/write requests to disk storage 405 in order to read or write data of the production file. For a first write to a file block since the last snapshot, the block mapping information for the file block must also be changed in the metadata of the file, since new data for the file block is written to a new storage block. The secondary storage server 402 not only changes the state of the file block from "READ-ONLY" to "READ-WRITE" in its local cache 408 but also changes the mapping information for the file block in the cache by removing the new mapping information for the file block from the write map 409 and writing it into the local cache 408. The secondary storage server 402 also sends a "Commit Metadata" request to the storage allocation server 403 so that the storage allocation server commits this change in mapping information to on-disk file metadata in disk storage 406.
In a preferred implementation, the disk storage 405 containing the file data is partitioned from the disk storage 406 containing the file metadata, and the storage server 402 does not have the right to directly access the file metadata in the disk storage 406. This helps to maintain the integrity of the file metadata.
As shown in
For example, the state and the storage location of each logical block in the range is returned in a list format, in which the state and the storage location is returned for the first block in the range, and for successive blocks in the range the block number, state, and storage location is returned only if there is a change in state or a change in storage location (other than a change in storage location offset not directly proportional to the offset of the block in the file). In particular, the list is a series of table entries ordered by the block offset in the file, and each entry includes a beginning block offset in the file, an ending block offset in the file, the state of the blocks from the beginning block offset to the ending block offset, and the address of the storage location for the beginning block offset for the file (or an indication that the beginning block offset is in a hole in the file). The secondary may cache the list by converting and linking the table entries into the index structure of the file system used in the secondary storage server, such as linking the table entries to the file inode or indirect block tree structure used in a Unix-based file system.
As shown in
In step 433, for each block that is not valid or valid but "READ-ONLY", then execution continues to step 435. In step 435, the storage allocation server checks its map of pre-allocated blocks. If the block is not already pre-allocated, then the storage allocation server allocates the block by selecting a free storage location for co-locality with neighboring blocks of the production file. Then in step 436, the storage allocation server returns to the secondary storage server an indication that the block is not valid and "READ-WRITE," and the storage location for the pre-allocated block, and processing of the "Allocate Space" request is finished with respect to the block.
For a file having a high degree of write activity, the storage allocation server can use a storage intensive method of pre-allocation in order to achieve co-locality of the pre-allocated storage with the shared storage having the same or neighboring block offsets in the file. For example, the storage allocation server could initially pair neighboring storage blocks having alternate sectors or tracks, and when a storage block is first committed, its pared storage block could be mapped to the same file and block offset in the file and placed in the storage allocation server's map 411 of pre-allocated blocks. Once a storage block becomes only in a snapshot (i.e., transitions to state S5 in
In step 447, if the block is not a hole in the file, then execution continues to step 450. In step 450, the secondary storage server reads data of the block from the block storage address for the block. Then in step 449, the secondary storage server returns the data of the block to the application or system process having requested the data.
In step 453, if the state of the block in the secondary storage server's local mapping cache is not "READ-WRITE," then the state is "READ-ONLY," and execution continues to step 457. In step 457, the secondary storage server looks up the block in its local write map of pre-allocated blocks for the file. Execution continues from step 457 to step 458 of
In step 459, the secondary storage server reads the storage address of the block from the write map. In step 460, the secondary storage server logs an intent to write data for a first write to the block since the snapshot. The log entry includes the offset of the block in the file and its new storage address from the write map, in order to assist recovery in the event of a system crash. Then in step 461 the secondary storage server begins writing the new data of the block to the storage address for the block, and also sends a "commit metadata" request to the storage allocation server. In step 462, the secondary storage server suspends processing of the write operation, and resumes upon completion of the write of the new data of the bock to storage. In step 463, the storage allocation server removes the new block from the write map, and in the secondary storage server's local cache of production file mapping information, replaces the storage address of the old block with the storage address of the new block and sets the new block state to "READ-WRITE." Finally, in step 464, the secondary storage server returns a write completed acknowledgement to the application or system process that requested the write operation.
Because the "intent to write" has been logged in step 460, it is not necessary for the secondary storage server to wait for the storage allocation server to commit the new metadata before returning a "write completed" acknowledgement to the application or system process that requested the write operation. Moreover, for more efficient transmission of requests to the storage allocation server, the secondary storage server can bunch up multiple "commit metadata" requests and send them all together to the storage allocation server. However, the secondary storage server should wait for the storage allocation server to complete all outstanding "commit metadata" requests for an application or system process when the application or system process requests the file to be closed or requests the "in cache" file system to be flushed to storage.
In step 452 of
In step 467 of
In step 458 of
In step 482, the storage allocation server unlinks any "READ-ONLY" blocks in the block offset range for the file from the production file mapping. These "READ-ONLY" blocks remain linked to the most recent snapshot file mapping so that they become owned by the most recent snapshot. Also, the storage allocation server unlinks the pre-allocated blocks in the block offset range in the file from the storage allocation server's local map of pre-allocated blocks, and links them to the production file mapping as committed "READ-WRITE" blocks (i.e., owned by the production file). In step 483, the storage allocation server commits the production file and pre-allocated block metadata changes to the file metadata in storage (406 in
Once the secondary storage server has completed the initial step 491, execution continues to step 492. In step 492 the secondary storage server and the storage allocation server set all blocks in their respective mapping caches for the production file to the "READ-ONLY" state. Then in step 493 the storage allocation server saves the current snapshot block map for the current snapshot of the production file and begins a new snapshot block map for the production file. For example, the storage allocation server clones the production file inode to create a new snapshot inode, and then changes a pointer to substitute the new snapshot inode for the current snapshot inode. Next, in step 494, the storage allocation server may request the secondary to release pre-allocated blocks or receive new mapping information for pre-allocated blocks. After any requested changes have been made to the maps of pre-allocated blocks in the secondary storage server and in the storage allocation server, execution continues to step 495. In step 495, the secondary storage server resumes processing of write operations upon the production file.
The storage block mapping protocol as described above can also be used between a storage allocation server and a storage allocation client in a cached disk array in order to provide virtualization of logical units of the storage in the cached disk array.
A major problem for customers of storage is managing the proliferation of storage arrays within their environment. Typically, each storage array manages its captive disks independently of any other storage array, and arrays from different vendors often provide incompatible management software and features. Virtualization of storage arrays is an approach that potentially allows customers to get a better return on their storage investment. A virtualization solution builds logical storage devices out of collections or fragments of physical storage devices. Inserting a logical-to-physical storage mapping layer in the I/O stack enables advanced storage features such as sparse provisioning, snapshots, and logical units (LUs) which span multiple physical arrays.
A cost associated with virtualization is the overhead of resolving the logical-to-physical address mapping for each I/O to a virtual device. One typical storage virtualization approach is to centralize this mapping calculation at a network switch blade which is positioned between the hosts and the storage array on the I/O path. Centralized map resolution simplifies the virtualization problem since it avoids the need for a map consistency and locking mechanism. However, the scalability of centralized map resolution is clearly a problem, since the only way to scale is to increase the power and cost of the mapping switch blade.
A solution to the problem of centralized map resolution is to put mapping for block resolution of virtual logical units (LUs) into a front-end director or a disk director of a cached disk storage array. A scalable solution results by putting space allocation and management of the mapping information under control of one or more storage allocation servers, and using a storage block mapping protocol to distribute the space allocation and mapping information from the storage allocation servers to front-end directors or disk directors in the data paths to the storage devices.
In a preferred implementation, the mapping for block resolution of a virtual LU is provided by a file system data structure and its mechanism for linking file system data blocks to the file system data structure. Thus, a virtual LU can be exported from the array as a file in the file system. In the array, however, a storage allocation client performs the logical-to-physical mapping in response to storage access requests from hosts in order to directly access the physical storage. As needed, the storage allocation client requests storage block allocations from a storage allocation server assigned to manage the mapping information for the LU, and once data is first written to the storage block allocations, the storage allocation client requests the storage allocation server to commit the storage allocations.
In a preferred implementation, the storage allocation client and the storage allocation server may function as described above with respect to the file server 402 and storage allocation server 403 in
In a preferred implementation, storage is virtualized by embedding software for the storage allocation server and the storage allocation clients into the front-end directors or disk directors of a cached disk array. This provides immediate benefits to the cached disk array by enabling new functionality. Furthermore, since the cached disk array has features not generally available on typical host platforms (i.e., large memory cache and non-volatile RAM), this embedding makes certain aspects of the storage virtualization more efficient.
As shown in
As shown in
Each of the disk directors 505 is programmed with storage allocation client software 508 for direct access to the disk drives in the disk array that are strung to the disk director. The storage allocation clients access client logs 511 of intent to write information in the global cache memory 504. These intent logs are significantly faster because they are held in protected global cache memory.
Incorporation of the storage allocation servers and the storage allocation clients into a cached disk array provide a number of additional advantages, including thin provisioning of the LUs, transparent relocation of the LUs, and scalable write-anywhere snapshots of the LUs.
Since UxFS files are sparse by default, mapping virtual LUs to files allows the presentation of a LU of a certain size, but which only consumes a subset of that amount of storage, namely those blocks which actually contain data written from the host. This feature is achieved by allocating space to the LU on a small extent granularity as writes are applied to the LU. The pool of storage that can ultimately be allocated to a sparse LU is a single UxFS file system. Multiple virtual LUs can be hosted on a single UxFS file system (as long as there are inodes available) and draw blocks of storage from the same pool. Given application level knowledge about which blocks are no longer used by the host, those blocks can be freed from one LU and made available for use by other LUs, simply by removing the blocks from the LU's container file mapping.
Virtual LUs can be transparently relocated. The storage block mapping protocol allows the storage allocation server to recall (invalidate) mapping information from storage allocation clients. This feature can be used to recall and update the mapping for a virtual LU when its data is moved from one physical location to another.
UxFS files can have a virtually unlimited number of snapshot copies (again, the only limit is the number of inodes available in the file system). Each snapshot is represented by a separate file system inode, with an independent block mapping, but data blocks which are the same in two successive versions are shared. This means that the various snapshots of a logical LU are stored compactly and efficiently, and the system has an efficient means to access the mapping for any snapshot. Maintenance of the snapshots minimizes copy-on-write by preserving existing data and allocating new space from the file system to hold new data. Copy-on-write is only desirable for partial blocks writes, when the I/O size is less than the file system allocation granularity (typically 8k bytes). Because the storage allocation clients are embedded in the storage array, many of these copy-on-write operations can be performed internally and thus very efficiently because they never cross the network.
It is also possible for a virtual LU to span multiple storage arrays, by building a single UxFS file system across physical devices contained in multiple storage arrays. In this case, the storage allocation client that performs the logical-to-physical mapping for the virtual LU is given access to all of the physical devices on which the virtual LU is built, for example, over a network link or a dedicated link between the multiple storage arrays. Specific examples are described below with reference to
There are several possible ways to embed storage allocation servers and storage allocation clients into a storage array. The choices revolve around the boundaries of each individual UxFS file system of LUs, and where the storage allocation server and storage allocation client software will be run. Each file system represents a shared pool of storage space from which virtual LUs (and their snapshots) can be provisioned. Thus, the storage allocation server is given access to disk drives containing the storage for metadata of the LU file system, and the LU file is given access to disk drives containing storage for data of the LU file system. This means that the intended scope of each pool of storage dictates where the storage allocation server and storage allocation clients for each pool should be run.
The simplest case is to bind a single pool from the set of disks managed by a single controller of the storage array. In this case, the storage allocation client can run on the controller itself The storage allocation server could also run on the controller, but since it is desirable for the storage allocation server to support many pools on many controllers, the storage allocation server would be better placed on a processor that has read/write access to as many controllers as possible. For this reason, in the cached disk array of
The data path of an I/O for the simplest case is shown in
The storage allocation client 508 is layered between the block I/O module 535 and the RAID/mirror module 537. The I/O requests pass from the block I/O module 535 to a file block mapping module of 534 of the storage allocation client, and from the file block mapping module 534 to the RAID/mirror module 537. The storage allocation client 508 has an FMP client module 533 for requesting block allocations and metadata from an FMP server module 531 in the storage allocation server 510. The storage allocation server 510 has a UxFS module 532 for managing the UxFS file system of the LU files.
In the simple case of
The next possible case is to allow a single UxFS pool to be spread over the disks of multiple controllers. Since the storage allocation client for the UxFS pool should have read/write access to all physical disks in the pool, it should only run on the disk director if there is a way to service read/write requests between different disk controllers. Otherwise the storage allocation client should run somewhere it can get read/write access to all the controllers on which the pool is built, such as on a front-end director.
As shown in
As shown in
In the configuration of
In step 643, to facilitate the separation of metadata from data and the pinning of the metadata in global protected memory, the UxFS on-disk format can be changed to group all of the file system metadata onto a contiguous set of volumes at the beginning of the file system address space. The remaining volumes which make up the file system are used for storing data only. Once the data and metadata in the file system are strictly partitioned, it becomes easy to pin the metadata volumes in global cache memory. This has the effect of improving the latency of metadata operations, and providing an extra level of protection for file system metadata.
In step 644, the file system metadata and transaction logs are fenced off from the storage allocation clients. Once metadata and data are strictly partitioned, the metadata volumes are made accessible only to the storage allocation server that owns the metadata, and not to any of the storage allocation clients. This significantly reduces the chance of write errors corrupting the file system metadata. A similar approach can be used for the file system transaction log. This log can be pre-pended to the same volumes which hold the file system metadata. Since only the storage allocation server needs access to the file system transaction log, the storage allocation clients can be fenced from the volumes holding it, increasing the safety of the log. Also, since the file system transaction log can be kept in fast global cache, the latency of logging transactions can be reduced.
Another powerful benefit of partitioning file system metadata from data is the ability to "upgrade" an existing array LU in place. By "upgrade," it is meant that it is possible to create a virtual LU which contains exactly the same data as an existing physical LU, with no data copy. In step 646, an existing physical LU is upgraded in place by linking the storage blocks to a new inode for a LU container file. Immediately after the upgrade, the virtual LU would contain the same physical blocks and data as the old physical LU. However, since the LU would now be virtual, it would be possible to start taking snapshot copies of the LU, and new writes to the LU would be able to consume space from the remainder of the UxFS pool. The reason this becomes possible is that once data and metadata within the file system are strictly partitioned, it becomes possible for a UxFS file system to dynamically consume contiguous ranges of physical blocks, regardless of size, simply by updating its volume mapping. Once the physical blocks of the LU are added to the pool of data blocks in the file system, a container file can be built which includes exactly the blocks of the physical LU.
The ability to upgrade an existing LU in place provides a smooth upgrade path for existing systems. Once an existing storage array is programmed with storage allocation and storage allocation clients, the system administrator can decide which existing LUs should be virtualized. The system administrator can then use the storage allocation server to build a UxFS pool encompassing those LUs, plus some extra storage to store snapshot copy information, build new virtual LUs which correspond to the old physical LUs, and finally, start taking snapshots of the new LUs.
Initially, when a first snapshot is taken, the space occupied by the old physical LUs cannot be reclaimed for general use by the pool, since the system has no knowledge about which blocks contain live data. However, as snapshots are taken, and the original "level 0" copy of the LU ages, at some point the system administrator may wish to refresh the oldest snapshot. At this point, all of the blocks of this snapshot that are not currently in use by newer versions (that is, all of the blocks which have been overwritten since the snapshot was created), are freed and returned to the pool for general use. In this way, the space of the original physical LUs can be reclaimed for general use as part of the natural snapshot lifecycle.
In step 647, the storage allocation server uses the FSCK utility for identifying potentially defective storage blocks. FSCK is the standard file system consistency checking tool for UxFS. In its current form it is designed to detect and repair certain file system data structure inconsistencies that can result from either software or hardware corruption. For instance, it can identify whether a single disk block is referenced by multiple files. In the storage array context, FSCK can be put to more general uses. One in particular is the identification and tracking of bad disk blocks. Currently, there is an absolute threshold of block failures that a single RAID group can tolerate before the entire RAID group is considered to be failed. FSCK can be enhanced to constantly check for failed blocks on disk, and when found, can take corrective action. In particular, the block can be marked as "used" within the UxFS pool. Therefore, in step 648, the storage allocation server excludes the potentially defective storage blocks from the LU storage pool, so that they will not be used again by any virtual LU. Also, the logical block of the virtual LU which was mapped to the failed block can be transparently remapped to a new location chosen from the pool. Thus a much larger number of failed blocks within a RAID group can be tolerated, since each failed block is isolated by the system and never used again. Continuing block failures simply result in a slow "leak" of available storage space.
In view of the above, there has been described a data processing system including a first data processor that pre-allocates data blocks of storage for use in a file system at a later time when a second data processor needs data blocks for extending the file system. The second data processor selectively maps the logical addresses of the pre-allocated blocks so that when the pre-allocated blocks are used in the file system, the layout of the file system on disk is improved to reduce block scatter and enhance I/O performance. The selected mapping can be done at a program layer between a conventional file system manager and a conventional logical volume layer so that there is no need to modify the data block mapping mechanism of the file system manager or the logical volume layer. The data blocks can be pre-allocated adaptively in accordance with the allocation history of the file system.
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Owner name: EMC CORPORATION, MASSACHUSETTS
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Dale Pattison - Relationship Assistant Manager at ANZ
Dale Pattison, Relationship Assistant Manager at ANZ from Melbourne Area, Australia Professional Experience,Educational Background,Friends,Photo,Email,Phone...
Racial Violence - Home - Contact Us
I would like to thank Professors Dale Pattison and Judith Stone of Western Michigan University and Professor Melvin Small of Wayne State University for helping ...
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classmates: Dale Pattison
Lewis & Clark High School, Spokane, WA, 1968-1972
classmates: Dale Pattison
Mission Secondary High School, Mission, BC, 1966-1970
classmates: Dale Pattison-ball
Montesano High School, Montesano, WA, 1993-1997
Bad news
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legacy.com: Dale Pattison
PATTISON, DR. DALE P. Dr. Dale P. Pattison passed away peacefully at home on April 16, 2012 following a lengthy illness. Dale was born September 6, 1928 in Kalamazoo to
findagrave: Pattison, Dale Philip
09/26/1928 - 04/16/2012, Kalamazoo (Kalamazoo County, Michigan)
Dale Pattison Kansas Death Records
Find death records for Dale Pattison in Kansas. A complete directory of death records and certificates organized by city, date, social security number, and more.
findagrave: Albert Dale Pattison ( - 1918) - Find A Grave Memorial
Birth: unknown. Death: Jul. 4, 1918. Note: Lieutenant, Australian Field Artillery. Age: 28. Burial: Boulogne Eastern Cemetery Boulogne-sur-Mer
Jessie Smith (Pattison) (1890 - 1969) - Genealogy - Geni
william pattison, walter pattison, goidie pattison. View the Record. Kate Jessie Dale Pattison in MyHeritage family trees (Gurman/Billings Tree) ...
Dale Pattison - Ancestry.com
Dale Pattison. ... All Death, Burial, Cemetery & Obituaries results for Dale Pattison. Search FiltersFilters. New. BroadExact. Dale. Pattison. Any: Mi. Year range: ...
Randy-J-Hicks - User Trees - Genealogy.com
Family Tree Maker user home page for Randy-J-Hicks.
John Dale Pattison - Ancestry.com.au
Name: Frederick John Dale Pattison. Residence: 1930 - city, North Sydney, New South Wales, Australia. Australia, Electoral Rolls, 1903-1980. Census & Voter ...
Download PDF - Project MUSE
Sites of Resistance: The Subversive Spaces of Their Eyes Were Watching God. Dale Pattison. Texas A & M University, Corpus Christi. Understanding the ...
Books & Literature
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Historian - Volume 57, Issue 4 - June 1995 - Wiley Online Library
... Frederic J. Baumgartner, Robert Weinberg, Gerard Clarfield, Albert A. Hayden, Marylou Ruud, Dale Pattison, D. R. Dorondo, Marjorie Morgan, John Beeler, ...
U.S. Supreme Court Transcript of Record Dale V. Pattison by U S...
9781270004424: The U.S. Supreme Court Transcript of Record Dale V. Pattison edition you want & more U S Supreme Court (Creator) new & used books ...
U.S. Supreme Court Transcript of Record Dale V. Pattison, ISBN...
Better Read Than Dead is a bookstore, a literary landmark that nourishes the neighbourhood's intellectual dynamics with regular author and community events.
Dale Pattison Collection - Google Books
Photocopies of a program book from the Ladies' Mutual Improvement Club of Oshtemo and Kalamazoo, Michigan, 1915-1916. Two Oshtemo Township, Kalamazoo County,...
Scientific Publications
Sites of Resistance: The Subversive Spaces of Their Eyes Were...
Dale Pattison. + Author Affiliations. Texas A & M University, Corpus Christi. email@example.com. Understanding the relationship between ...
Dale Pattison's (pattisondale) Pictures, Photos & Images | Photobucket
View and follow Dale Pattison (pattisondale)'s photos and videos on Photobucket .
Video & Audio
Dale Pattison - YouTube
Dale Pattison. Videos · Playlists · Channels · Discussion ... Uploads. Dale Pattison uploaded a video 7 hours ago. 4:30. Play next; Play now ...
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Student in court over killing - Sunderland Echo
Dale Pattison, 18, of Manila Street, Sunderland, appeared at Newcastle Crown Court yesterday in connection to the death of Raymond Quigley.
Heworth ARLC toppled by champions East Leeds (From York Press)
Heworth ARLC were undone by free-scoring winger Jason Priestley as they lost 34-10 at home to National Conference League division two champions…
School Zone blog: 57th anniversary of Brown v. Board of Education...
It took 17 years for Brown v. Board of Education to hit home in Kalamazoo. And then it hit with a vengeance
Dr. Dale Pattison - Western Michigan University
The Department of History is mourning the loss of WMU Graduate and Professor Emeritus Dr. Dale Pattison, who passed away at his home on April 16, 2012, ...
Reports & Statements
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Google Groups: Hitler as a Pagan?
29.07.1998: Dale Pattison mer...@charleston.quik.com alt archaeology The term pagan is a
pinterest.com: Dale Pattison on Pinterest
See what Dale Pattison (jdklp) has discovered on Pinterest, the world's biggest collection of everybody's favorite things.
Conference Schedule | SSAWW Conference 2012: Citizenship and Belonging
Dale Pattison (Arizona State University): "The Trauma of Erasure: Helena María Viramontes' Their Dogs Came With Them and the New Metropolis". ASLE, the ...
Google Groups: Printing Messages
22.01.1997: Dale Pattison d...@thinktech.com.au comp groupware groupwise One of my
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Dale Pattison's social stories · Storify
Cover photo for Dale Pattison. Dale Pattison. DalePattison. 0stories. 0followers · 1following. Parksville, British Columbia. https://storify.com/DalePattison.
DALE PATTISON (1926-2000) | Social Security Death Index | SSDI
Dale Pattison was born in 1926 and passed away in 2000 . Access the SSDI on Mocavo and find your ancestors for free!
Bruce Dale Pattison, lawyer in Johannesburg
Phone and address of Bruce Dale Pattison, Johannesburg (2192). Directory of lawyers in Johannesburg Metro.
Dale Pattison at Arizona State (ASU) | Uloop
Rate and Review Dale Pattison at Arizona State (ASU), along with other Professor Ratings in Tempe, Arizona.
Head To Head | Dale Pattison vs Rahim Esmail H2H
Head To Head. Dale Pattison vs Rahim Esmail all matches, with stats on their H2H rivalry. ATP & WTA Dale Pattison head to head tennis search.
Dale Pattison Accepts a Position at Texas A&M—Corpus Christi |...
Please join in congratulating Dale Pattison, PhD in Literature 2013. He has accepted a position as Visiting Assistant Professor at Texas ...
U.S. Supreme Court Transcript of Record Dale V. Pattison : ISBN...
Buy the U.S. Supreme Court Transcript of Record Dale V. Pattison with fast shipping and excellent Customer Service. Rakuten.com
Dale Pattison at Think Technology from Geelong, Australia
Find out more about Dale Pattison, Manager, Think Technology, on Tradeloop, the original and largest B2B Trading Network for the Secondary IT Market
Dale Pattison | Extended Profiles | MLA Commons
Dale Pattison. @pattisondale active 1 year, 1 month ago ... Name. Dale Pattison. Profile picture of Dale Pattison. Activity · Profile · View · Sites 0 · Contacts 0 ...
Rate Dale Pattison, a professor at Arizona State University.
Students can rate Dale Pattison, English professor at Arizona State University, to inform other students about their experiences with this college professor.
Norman Dale Pattison b. Private: Families of NW Pennsylvania
Norman Dale Pattison b. Private: Families of NW Pennsylvania
Dale Pattison Twitter Stats - Twitter Counter
Get a clear overview of Dale Pattison (@dalepatt)'s twitter stats or see the graph of your own Twitter account.
Dale Pattison | Draw Your Professor
See Drawings ofDale Pattison: professor of English at Arizona State U - Tempe!
Kate Jessie Dale PATTISON 1890–1969 - Mid North Coast Pioneers
Player Profile -
Information on every club in the National Conference League
Dale P. Pattison obituary | WMU News | Western Michigan University
Dr. Dale P. Pattison, a retired Western Michigan University faculty member, died April 16. He was 83. Pattison joined the WMU faculty in 1963 ...
Pattison, Dr. Dale
Dr. Dale P. Pattison. Born in Kalamazoo, Michigan on Sep. 6, 1928. Departed on Apr. 16, 2012 and resided in Kalamazoo, MI.
John Dale PATTISON –1938 - Mid North Coast Pioneers
Pattison vs Esmail H2H Stats - Betting Odds & Tips · MatchStat
Rahim Esmail vs Dale Pattison head to head stats and betting tips. Make a free bet now at the best odds with our virtual game & odds checker system. Rahim...
UVHL : Player : Dale Pattison
Games Played, Goals, Assists, Points, PIMs. 0, 0, 0, 0, 0. T-106th, T-81st, T-92nd, T-97th, th. NONE, Log in to update your UVHL Profile: My UVHL. Dale Pattison ...
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text:
Respond to Review of Green Vegetarian Cuisine at Alon
by ChibiArtist
Review: I currently live in LA but I'm originally from SA, so while visiting family for the holidays I ate at this restaurant. Within one week of being in SA I've been here twice. Love it! The atmosphere is really nice. The staff are great and friendly. The food is amazing. I had the roasted veggies bowl, sweet potato fries, sweet raspberry tea, and avocado eggs Benedict (made vegan). Everything was amazing, something I could of found in LA. I wanted to get a shirt but they didn't have smalls but the shirts are adorable. Some say "vegetarians are better lovers", "fake meat saves lives", "give peas a chance", etc. they also have cupcakes which are displayed in a glass case in the front of the store before even entering I had already decided I had to have one. They didn't disappoint. I had a chocolate cinnamon cupcake and it was excellent. Everything on they're menu can be made vegan but what I like the most is that all the desserts, breads, dressings, and sauces are vegan. No need to worry if there's milk or eggs. I will definitely come here every time I'm in town more than once lol.
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text:
Cartoon Network (ended 2003)
I'd be really interested to hear more about this, too! It'd be so awesome to see a Trigun movie.
My lord I wanna play that!
Where'd u hear it?
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text:
Let's be honest here. It doesn't take much to make burgers, but it takes a lot of skill to make a GOOD burger. So if you're opening a joint that only sells this year-round classic, then every component has got to be bloody fantastic.
I asked for my patty to be cooked medium, et voila! They did not disappoint.
You can wash it down with homemade lemonade (£2.50)
Or my favourite: scotch whiskey + ginger ale. ;)
Brixton Village, Unit 12, SW9 8PR London | +44 (0)20 7733 7963
Average spend pp: £15


text:
Erick Silva Newswire
Comprehensive Real-Time News Feed for Erick Silva.
Results 1 - 20 of 126 in Erick Silva
UFC Welterweight Rick Story Sidelined 3 Months Following Neck Surgery w/Photo
... his neck. Story is hoping to return to the Octagon sometime in early 2016. "The Horror" was scheduled to face Erick Silva at UFC Fight Night in Saskatchewan on Aug. 23, but the aforementioned injury forced him to withdraw from the card. Prior to ... Comment?
Silva hints at injury for 'ridiculous' performance against Magny
Welterweight slugger Erick Silva hit another roadblock on Sunday Night at UFC Fight Night 74. The Brazilian was expected to cruise past short notice replacement Neil Magny in the co-main event but found himself on the losing end of a lacklustre split decision. Silva, known for his devastating power and savvy submission game, never seemed to find his groove in the Octagon against Magny, as the American stayed out of harms way and utilized his range to outland his opponent. Comment?
Erick Silva embarrassed by his body, suggests injury contributed to loss in Saskatoon
The Brazilian welterweight lost a split decision to Neil Magny at the UFC Fight Night 74 mixed martial arts event last Sunday night in Saskatoon , after turning in a performance that was far below what "Tiger" fans are used to seeing. "Honestly, I did not like my performance, and was overcome, I was embarrassed by my performance, my body, I know that I am much more than that Erick who presented yesterday, my main mistake was insisting on fighting, and put not only my 'head' in play as my coaches and my doctors did their best as far as possible so that I at least was the fight. Comment?
UFC Fight Night 74 '6th Round' post-fight show for 'Holloway vs. Oliveira' in Saskatoon
... took place inside SaskTel Centre in Saskatoon, Saskatchewan, Canada. Neil Magny winning a split decision over Erick Silva after taking the fight on two-weeks notice. The back-and-forth, "Fight of the Night" winner between Patrick Cote and Josh ... Comment?
Erick Silva 'embarrassed', critical of his 'ridiculous performance' against Neil Magny
Erick Silva had the chance to win his third straight in the UFC Octagon for the first time Sunday night, but came up short against Neil Magny in the co-main event of UFC Fight Night 74 in Saskatoon, Canada. Disappointed with his performance after a split decision loss, Silva took on his social media to apologize, and hinted at a possible injury he suffered before the fight. Comment?
The Walkmen: UFC Fight Night 'Holloway vs. Oliveira' Walkout Songs w/Photo
... through the arena. He clearly did not regret taking the co-main event slot opposite former Jungle Fight champion Erick Silva on short notice, just 22 days after his lopsided loss to 2007 Abu Dhabi Combat Club Submission Wrestling World Championships ... Comment?
Neil Magny vs. Erick Silva full fight video highlights from UFC Fight Night 74 in Saskatoon
Top welterweight standout Neil Magny got back into the win column by turning away Erick Silva in the co-main event at UFC Fight Night 74 last Sunday night inside SaskTel Centre in Saskatoon, Saskatchewan, Canada, despite taking the fight on just two weeks notice. The split decision win is likely to help Magny forget about his submission loss to Demian Maia earlier this month in Brazil, one that snapped a seven-fight winning streak and sent the 28-year-old prospect tumbling to No. Comment?
Top 5 Fights To Make In The Wake Of UFC Fight Night 74 w/Photo
... made a successful Octagon debut against Angela Magana. Magny made the most of his short-notice clash against Erick Silva by producing a split-decision victory. He dominated the Brazilian with his striking volume and got back into the win column ... Comment?
The Good, Bad and Strange from Fight Night 74 w/Photo
... interested in wasting any time getting things back on track. When Rick Story was forced out of his bout with Erick Silva at Fight Night 74, the Colorado-based fighter jumped at the opportunity to face the Brazilian striker. While he was the underdog ... Comment?
Holloway takes down Oliveira through TKO at UFC Fight Night Saskatoon w/Photo
... can't fight, you can't pay your bills. It is sad to see."\ Neil Magny (16-5) won a split decision over Brazilian Erick Silva (18-6) in the night's co-main event. The fight had action in the first round but became a stand-off boxing match for the ... Comment?
UFC Saskatoon results: Neil Magny beats Erick Silva by split decision
Check out the result of the UFC Fight Night: Holloway vs. Oliveira co-main event between welterweights Neil Magny and Erick Silva. The UFC Fight Night: Holloway vs. Oliveira co-main event was supposed to be Erick Silva against Rick Story , but when Story pulled out, Neil Magny stepped in on short notice despite coming off a submission loss to Demian Maia on August 1st. Comment?
UFC Fight Night: Holloway vs. Oliveira results and post-fight analysis
... fight one day, maybe even as early as December, I highly doubt you'll ever see it in a stadium in Hawaii. Erick Silva looked different. I'll let you draw your own conclusions as to why that may be. He was the same fighter though - firing away in ... Comment?
Charles Oliveira Injury Gives Max Holloway Victory in UFC Fight Night Headliner w/Photo
... Magny rebounded from his first loss in nearly two years with a split decision over former Jungle Fight champion Erick Silva . Two of the three cageside judges scored it 29-28 and 30-27 for Magny, while a third cast a dissenting 29-28 nod for Silva. ... Comment?
UFC Fight Night 74: Holloway vs. Oliveira Post Fight Press Conference w/Photo
UFC Fight Night 74: Holloway vs. Oliveira is in the books, and it was Max Holloway who walked out of the main event when Charles Oliveira stopped the fight with an injury early in the first round. In the co main event, Neil Magny was able to come in on late notice to take a Split decision win over Erick Silva. Comment?
UFC Fight Night 74 results: Neil Magny dominates Erick Silva, somehow ...
The three judges scored it 29-28, 28-29, and 30-27. Magny, 28, ended up on top of Silva early on in the first after the Brazilian missed on a poorly attempted spinning-back fist. Comment?
UFC Fight Night 74 Results: Neil Magny Decisions Erick Silva in Co-Main Event w/Photo
... a risky move for an athlete; the decision can make or break a fighter's career. Dangerous Brazilian welterweight Erick Silva was originally scheduled to face Rick Story at UFC Fight Night 74 on Saturday, but an injury eleven days ago forced Story ... Comment?
Top 10 Pound-For-Pound Boxers Today w/Photo
... oway looms as a fan-friendly and important main event at UFC Fight Night 74. Read More Exciting UFC welterweight Erick Silva can prove he's top-15 material by getting past Neil Magny this weekend at Fight Night 74. Read More Comment?
UFC Fight Night: Holloway vs. Oliveira staff picks and predictions
... Max Holloway and Charles Oliveira. Holloway is favored by the Bloody Elbow staff to get the win, as is Erick Silva over Neil Magny in the co-main. I'm personally glad that the entire staff did go with Nikita Krylov to win his fight on the prelims. ... Comment?
UFC Fight Night 74: Erick Silva Needs To Keep Positive Momentum Rolling Against Neil Magny w/Photo
Talented Brazilian welterweight Erick Silva again has the chance to finally take his career to the next level when he tangles with Neil Magny this weekend at UFC Fight Night 74 in Saskatoon. Magny accepted the short-notice assignment against the explosive Brazilian after Rick Story was forced out of the co-main event due to injury. Comment?
Erick Silva: Neil Magny hasn't fought someone as tough as me before
Erick Silva 's return was moved from June to August, and his opponent switched due to an injury, but the Brazilian welterweight is finally ready to step inside the Octagon. Initially slated to meet Rick Story on June 26, and then moved to Aug. 23 due to visa issues, Silva takes on Neil Magny after Story suffered an injury, and the constant changes forced the Brazilian to adapt in training ahead of Sunday's UFC Fight Night 74 co-main event bout. Comment?
Copyright © 2015 Topix LLC


text:
Release Date:
November 19, 2012
Culled from singles that date back to 2001, a track from the EP Newborn ("None One"), and B-sides from 2011's well-received Build a Rocket Boys, Dead in the Boot offers up a rare glimpse into the often insular song-building world of the Manchester outfit. Decidedly less bombastic and ornate than the majority of the band's more anthemic album offerings, Dead in the Boot is a quieter, more abstract affair that feels surprisingly autonomous. Elbow have always straddled the line between stadium-ready house band and a band that just wants to stay in the house, lock the doors, and be left alone to die, and it's the latter persuasion that informs the majority of the collection's 13 cuts. The brittle "Whisper Grass" leads things off in a decidedly mellow direction, but like many of the band's best moments, it recoils and strikes when the listener least expects it, spewing a sinewy stream of distorted venom around the two-minute mark that changes the whole timber of the song. Likewise, the brooding "McGreggor" coolly struts its bluesy swagger before allowing Guy Garvey to indulge in some "Grounds for Divorce"-inspired caterwauling, probably due in large part to the track's live setting. That said, the rest of Dead in the Boot barely moves the VU meter, choosing instead to occupy that creepy corner of the room where doubt, fear, shame, indulgence, and misplaced rage go to fester when the clock sneaks past two in the morning, the rain flirts with ice, and there's only a single, half-wet cigarette left in the pack. [A Deluxe Edition featured bonus tracks.] James Christopher Monger, Rovi
- Genre: Rock
- Category Britpop, Dream Pop
- Artist: Elbow
other Info.
- Online Item #: 14267889
- Store Item Number (DPCI): 244-25-0709
- Made in the USA or Imported