Linux Buffer Cache and Block Device

Linux Content Index

File System Architecture – Part I
File System Architecture– Part II
File System Write
Buffer Cache
Storage Cache

Working on a flash file system port to Linux kernel mandates a good understanding of its buffer management, sometimes the excavation of internals can be a bit taxing but it’s needed for executing a well integrated port. Hopefully this write up will help someone who is looking to get a good design overview of the kernel buffer/page cache.

VFS work on top of pages and it allocates memory with a minimum granularity of the page size (usually 4K) but a storage device is capable of an I/O size much smaller than that (typically 512 bytes).  A case where the device write granularity is set to page size can result in redundant writes, for example if an application modifies 256 bytes of a file then the corresponding page in cache is marked as dirty & flushing of this file results in a transfer of 4K  bytes because there is no mechanism for the cache to identify the dirty sector within this dirty PAGE.

Buffer Cache bridges the page and the block universes, a page associated with a file  will be divided into buffers & when a page is modified only the corresponding buffer is set as dirty, so while flushing only the dirty buffer gets written. In other words a buffer cache is an abstraction done on top of a page, its simply a different facet to the same RAM memory. The data structure called “buffer_head” is used for managing these buffers, the below diagram is a rather crass depiction of this system.

Buffer Cache can traverse the list of these buffers associated with a page via the “buffer_head” linked list and identify their status; note that the above diagram assumes a buffer size of 1K.

How about streaming writes?

When dealing with huge quantities of data, the method of sequential transfer of small sectors can lead to a pronounced overall read-write head positioning delay and there is also this additional overhead due to multiple set up and tear down cycles of transfers.

Lets consider an example of a 16K buffer write operation consisting of four 4K pages, the memory mapping for these 4 pages are given in the table below. For example, sectors 0 to 3 of a file is located in the RAM address 0x1000 to 0x1FFF, and it’s mapped to storage blocks 10, 11, 12 & 13.

Blindly splitting the 16K into 16 buffers and writing them one by one will result in as many transfer set ups and teardowns, and also an equivalent number of read-write head position delays.

How about we reorder the writing of pages in the following manner?

Page1 >>> Page 3 >>> Page 2 >>> Page 4

The above sequence is the result of sorting the pages in an increasing order of block mappings, ie: from 10 till 25, a closer look will also reveal that the adjacent transfers have addresses running from 0x1000 till 0x4FFF and hence contiguous in RAM memory. This results in an uninterrupted  transfer of 16K bytes, which seems to be the best case optimization.

How does Linux achieve this gathering?

Let’s try to comprehend how bock driver subsystem skews application’s order of writes to achieve the previously described optimization.

Algorithm implemented by the above architecture is explained below

Step 1: A Buffer is encapsulated in a BIO structure.
File system buffers are encapsulated in a BIO structure  (submit_bh function can give the details). At this stage all the 16 buffers in our example are wrapped inside its corresponding BIO struct. We can see the mapping in the below table.

Step 2: A BIO is plugged into a request.
BIOs are sequentially submitted to the below layers and the those mapped to adjacent blocks are clubbed into a single “request“.  With the addition of the request struct we have the below updated mapping table, as you can see “Request 0” combines BIO 0,1,2 & 3.

Prior to adding the requests to the global block driver queue there is an intermediate task_struct queue, this is an obvious optimization aimed at clubbing the adjacent BIOs into one request and to minimize contention for the block driver queue. Yes, a global queuing operation cannot happen in parallel to another queuing or a de-queuing.

Step 3: Merge requests into the global queue.
Merging a request into the global queue will reorder and coalesce the adjacently mapped “requests” into one and in our case we will finally have only one single request which will envelop all the BIOs.

Step 4:Finally we need to identify the contiguous memory segments.
Block device driver will identify consecutive memory segments across adjacent BIOs and club them before initiating a transfer, this will reduces the I/O setup – tear down cycles. So we will have a segment attribute added to our final table which is given below.

The block device driver can also have a request scheduler, this scheduler algorithm will prioritize and pick requests to be de-queued from the global queue for I/O processing. Please note that the steps for read requests also follow the above framework.

Inference

The above bells and whistles are irrelevant when it comes to flash memories, so avoiding all this baggage might make a lot of sense.

Linux File System Stack – 2

Linux Content Index

File System Architecture – Part I
File System Architecture– Part II
File System Write
Buffer Cache
Storage Cache

A Linux file system is expected to handle two types of data structure species — Dentries & Inodes. They are indeed the defining characteristic of a file system running inside the Linux kernel. For example a path “/bm/celtic” contains three elements, “/” , “bm” & “celtic”, so each will have its own own dentry and inode. Among a lot of other information a dentry encapsulates the name, a pointer to the parent dentry and also a pointer to the corresponding inode.

What happens when we type “cd /bm/celtic”?

Setting the current working directory involves pointing the process “task_struct” to the dentry associated with “celtic”, locating that particular entry involves the following steps.

  1. “/” at the beginning of the string indicates root
  2. Root dentry is furnished during file system mount, so VFS has a point where it can start its search for a file or a directory.
  3. A file system module is expected to have the capability to search for a child when the parent dentry is provided to it. So VFS will request the dentry for “bm” by providing its parent dentry (root).
  4. It’s up to the file system module to find the child entry using the parent dentry. Note that the parent Dentry also has a pointer to its own inode which might hold the key.

The above sequence of steps will be repeated recursively. This time the parent will be  “bm” and “celtic” will be the child, in this manner VFS will generate a list of Dentries associated with a path.

Linux is geared to run on sluggish hard disks supported with relatively large DRAM memories. This might mean that there is this ocean of Dentries and Inodes cached in RAM & whenever a cache miss is encountered, VFS tries to access it using the above steps by calling the file system module specific “look_up” function.

Fundamentally a file system module is only expected to work on top of inodes, Linux will request operations like creation and deletion of inodes, look up of inodes, linking of inodes, allocation of storage blocks for inodes etc.

Parsing of paths, control cache management are all abstracted in kernel as part of VFS and buffer management as part of block driver framework.

How about writing to new file?

  1. User space communicates the buffer to be written using the “write” system call.
  2. VFS then allocates a kernel page and associates that with the write offset in the “address_space” of that inode, each inode has its own address_space indexed by the file offset.
  3. Every write needs to eventually end up in the storage device so the new page in the RAM cache will have to be mapped to a block in the storage device. For this VFS calls the “get_block” interface of a the file system module, which establishes this mapping.
  4. A copy_from_user_space routine moves the user contents into that kernel page and marks it as dirty.
  5. Finally the control returns to the application.

Overwriting contents of a file differ in two aspects – one is that the offset being written to might already have a page allocated in the cache and the other is that it would be already mapped to a block in the storage. So it’s just a matter of memcpy from user space to kernel space buffer. All the dirty pages are written when the kernel flusher threads kick in, and at this point the already established storage mapping will help the kernel identify to which storage block the page must go.

Reading a new file follows the similar steps but it’s just that the contents needs to be read from the device into the page and then into the user space buffer. If an updated page is encountered, the read from storage device read is of course avoided.

Linux Kernel Caller ID

Kernel Caller ID

Linux kernel boasts of a useful debug mechanism for printing the caller of a function.

printk(“Caller is %pS\n”, __builtin_return_address(0));

Without a JTAG debugger this is my primary tool to figure out the call stack. (Any better ideas?). Took couple of hours to dissect this mechanism but it was worth it.

There are two main parts to this:
1. Get the caller address.
&
2. Map the caller address to the caller name.

Step 1 : __builtin_return_address(0)
ARM assembly clearly shows that this is an assembler directive which simply fetches the caller address from the stack (or the return address register perhaps?)

Step 2: How do we map this address into a symbol string?
In simple words, when kernel is linked it created a compressed version of symbol table and parsing this provides the mapping from address to string. A one to one mapping of address and the ASCII string would lead to a gargantuan sized binary so we need some form of compression.

So the algorithm tends to encodes the symbol by exploiting the repeating character patterns, for example a string “__exit” might be represented by an arbitrary code 0x34, so the idea is to identify the these patterns and generates a custom representation of a string. Elegant and effective!! This works just fine on loadable kernel modules also, because the dynamic loading process takes care of this. More details might need a look into insmod!

Please look into the C file linux/kernel/kallsyms.c for discovering more.

Linux File System Stack – 1

Linux Content Index

File System Architecture – Part I
File System Architecture– Part II
File System Write
Buffer Cache
Storage Cache


Why File System as a Loadable Kernel Module(LKM) ?

As you can see, user space idea didn’t pan out quite well:http://tekrants.me/2012/05/22/fuse-file-system-port-for-embedded-linux/

VFS Data Structures

Inode is  probably the most critical abstraction which defines a VFS file entry — This represents every file/directory/link within a file system. If your file system is like FAT and lacks a clear “inode”, then a translation layer will be needed. Eventually it’s about extracting the file information associated with a Linux inode from the file system specific data structures.

File System Initialization Sequence

1. Register file system mount & unmount call backs with the VFS.
2. Mount call is responsible for creation and registration of a root directory inode .
3. Root directory inode is essentially the point of entry to the volume. It furnishes specific function pointers later invoked by VFS for inode related operations (like create) and file operations (like open,read) & directory operations (like readdir).

The above three steps and your file system module is all set, this means Linux will have enough information to translate an “Open” call from application to the file system specific internal open call, thanks to the function pointers inside the root inode.

Dentries

Another kernel structure which exists for every file & directory is a Dentry. For example, accessing a path “/mnt/ramfs” will lead to creation of two in memory dentry structures. One each for “mnt” and “ramfs”. Note that “ramfs” dentry will have a parent pointer to “mnt” dentry and a pointer to its own VFS inode. A Dentry in fact encompasses the attributes like name & handle to parent directory of a file system entry. One of the rationales behind separation of an Inode from these attributes is the existence of file links, where a single Inode is shared across multiple Dentries.

Opening a file — Easily said than done!

a. Consider opening a file with the path “/mnt/ramfs/dir1/dir2/foo.txt”
b. The dentry elements in the above path are “mnt”, “ramfs”, “dir1”, “dir2” & “foo.txt”
c. “mnt” dentry will be part of Linux root file system, all dentries are part of a hash table, the string “mnt” will map to a hash table entry giving its dentry pointer. VFS gets the inode pointer from this dentry and every directory inode has a callback function for look-up operation listing on its file/directory entries.
d. Look up called on “mnt” inode will return the inode for “ramfs” along with its dentry.
c. This is an iterative process and eventually VFS will figure out the inodes & dentries of all the elements in a file path.
d. Note that the Inode associated with “foo.txt” will give the open function pointer to invoke the open call specific to the file system driver.

VFS

A file system ported to Linux is expected to populate the fields of VFS data structures like Inodes and Dentries so that Linux can understand and convey the file attributes and contents to the user. The obvious differentiating factor across file systems like ext4, UBIFS, JFFS2 etc are their respective algorithms, which also defines the internal data structures and device access patterns.

How Dentries/Inodes are represented and accessed from a storage is specific to file system and this inherently defines their strengths and weaknesses. In crude terms, a file system in Linux comprises of a set of call backs for managing generic VFS data structures, basically the Inodes, Dentries, file handlers etc. So we have inode data structure and corresponding associated inode operations, we have file pointer data structure and file operations, dentry data structure and dentry operations and so on.

The crux of a Linux file system is its ability to talk in Linux kernel language of Inodes and Dentries. Also, unless it’s a read only volume this interpretations needs to be done in reverse too. When user makes changes to a file then a file system needs to comprehend the Linux talk and translate those changes into a representation which it might have on the storage. Undoubtedly, comprehending Linux VFS mandates deep understanding of Kernel data structures which might mean that a file system writer needs to have a kernel specific layer in the file system code, this undesirable complexity can be immensely reduced  by the use of kernel library functions.

Functions which usually start with “generic_” can be interpreted as such a helper function which abstracts the kernel specifics from a kernel module, this is widely used for file system operations like “read”, “write” and even unmount. The usage of generic helper functions within a kernel module can be confusing when studying the code, because they tend to blur the kernel and a module specific boundaries, this overlap is convoluted but an extremely effective way to avoid kernel dependencies.

Image Source : http://wiki.osdev.org/images/e/e5/Vfs_diagram.png

Some design philosophy

The design thought behind Linux VFS seems to of “choice”, other than a bare minimum there is always a choice provided to the kernel hacker regarding the implementation of an interface. He or she can either use a generic function, create a custom one or simply set it to NULL. This is an ideal approach for supporting a plethora of widely differing devices and file systems, quite easily visible when we look at the code of an ext4 where there is buffer cache abstraction usage over page cache, compared to page cache sans buffer for UBIFS, versus a direct to the device approach of JFFS2. Accommodating all these widely varying designs requires a flexible rule of law driven framework where everyone can find their niche and yet not impinge on the functionality of another kernel module.

A Linux File System – 2

FUSE File System Performance on Embedded Linux

We ported and benchmarked a flash file system to Linux running on an ARM board. Porting was done via FUSE, a user space file system mechanism where the file system module itself runs as a process inside Linux. The file I/O calls from other processes are eventually routed to the FUSE process via inter process communication. This IPC is enabled by a low level FUSE driver running in the kernel.

http://en.wikipedia.org/wiki/Filesystem_in_Userspace

 

 

The above diagram provides an overview of FUSE architecture. The ported file system was proprietary and was not meant to be open sourced, from this perspective file system as a user space library made a lot of sense.

Primary bottleneck with FUSE is its performance. The control path timing for a 2K byte file read use-case is elaborated below. Please note that the 2K corresponds to NAND page size.

1. User space app to kernel FUSE driver switch. – 15 uS

2. Kernel space FUSE to user space FUSE library process context switch. – 1 to15 mS

3. Switch back into kernel mode for flash device driver access – NAND MTD driver overhead without including device delay is in uS.

4. Kernel to FUSE with the data read from flash – 350uS (NAND dependent) + 15uS + 15uS (Kernel to user mode switch and back)

5. From FUSE library back to FUSE kernel driver process context switch. – 1 to 15mS

6. Finally from FUSE kernel driver to the application with the data – 15 uS

As you can see, the two process context switches takes time in terms of Milliseconds, which kills the whole idea. If performance is a crucial, then profile the context switch overhead of an operating system before attempting a FUSE port. Seems loadable kernel module approach would be the best alternative.

Pattern Recognition

The term “pattern recognition” always comes up when someone is discussing image mining, handwriting recognition and the related spaces. Basically, whenever there is an explicit requirement for discovering meaningful information after casting out the noise.

Recently ran into this word biomimesis, study of biological systems for engineering solutions for our “real world” problems. This seems to be an interesting discipline of searching for problems which are already solved in nature and then adapting it for our purpose! Prima facie this might seem trivial but imagine a case where a Morpho butterfly wings has inspired creation of MIRASOL display technology for consumer electronics. Now imagine how arduous it could be to explain the challenges involved in creating low power electronic displays to a biologist!

How can problem solving methods be abstracted enough to be applied across a myriad of domains? The big challenge here may not be bringing together a bunch of experts from widely different domains but it might be about successfully harnessing their understanding and also in finding ways to make them all talk in the same “language”.

Here is another article on a similar topic. Mentions how the guys at Wall Street and also software giants like IBM hire astrophysicists and theoretical mathematicians. When your area of expertise mandates abstract thinking, solving software technology problems and identifying patterns across financial data might be a walk in the park.

In NAND memory, why is there a limit to the times we can make a partial write to a page?

A question posted on Quora — “In NAND memory, why is there a limit to the times we can make a partial write to a page? Is it because of write disturb. If so how does it happen? Does the same apply to partial reads?”

Having worked for some time on SLC NANDs, few points which comes to my mind are :

NAND data sheets from Samsung, Toshiba & Micron tend to limit the number of times a page can be partially programmed (usually to 4), this may be a circuit design decision which might have aided in some optimization with cost/speed/reliability etc. The exact reason could be only given by the NAND chipset designers.

Did observe that if we try, then we could partially program pages for more than the actually mentioned 4 times. But reliability is a question here! Four cycles of partial programming is a good number, if your software is doing frequent partial writes then it is inviting a performance hit.

Read disturbances are known to cause bit flips and excessive number of block erase cycles can lead to bad blocks. Not aware of any “write disturb”on SLCs.

NAND page read protocol will always fetch the data from the cells into an internal NAND chip RAM buffer from there the data can be sampled though any multi pin interface. In that sense, have not come across any hardware which allows partial read of physical NAND pages. We can definitely do random partial reads on a page after it is fetched from the NAND cells into this internal NAND chip buffer, but then this can be executed unlimited number of times

Relevance of Quad Core processors in mobile computing

October edition of EE Times Europe has an article written by +Marco Cornero from ST Ericsson explaining how Quad core processors for mobile computing is ahead of its time. The following tweet was send from the official ST Ericsson account.

“Quad cores in mobile platforms: Is the time right?” An article in EETimes Europe written by Marco Cornero, ST-Ericsson http://ow.ly/6TrDo

Please note that you might need to log in to EE Times website to access the full article along with the whole October edition.

In some ways this is quite a brazen stance taken by ST Ericsson.

1. +Marco Cornero states that there is a 25% to 30% performance overhead on each core while moving from Dual to Quad, this is due to systemic latency attributed to L1/L2 cache synchronization, bus access etc.

2. This overhead will mandate that for a Quad core to out perform Dual core each software application needs to have 70% of its code capable of executing in parallel. How this is calculated is clearly explained in the article.

The article also argues that there are not many complex use cases which can create multi-tasking scenarios where there is optimal usage of all the four cores and multiple other arguments which seems to conclusively prove that Quad cores is a case of “Diminishing returns” (See below quote from ST Ericsson CEO Mr.Gilles Delfassy)

“We aim to be leaders in apps processors, but there is a big debate whether quad core is a case of diminishing returns,” – Gilles Delfassy

The whole Article hinges on one fact that there will be a 25%-30% performance overhead on each core while moving from Dual to Quad, but isn’t this purely dependent on hardware? This figure may hold good for ST Ericsson chip-set, but what about ASICs from Qualcomm, NVIDIA, Marvel etc?

The crux of the argument is this very overhead percentage and if we bring this overhead marginally down to 20-25%, then only 50-60% of the application code needs to execute in parallel for optimal usage of Quad core. This situation is not so bad. Is this inference inherently flawed? The fact that Qualcomm and NVIDIA are close to bringing out Quad core solutions to market makes me wonder about ST Ericsson’s claims!

Stack Memory Usage

Real Time Embedded software running on ASICs with limited memory and capabilities tend to lack paging and virtual memory units also. Which in turns places constraints on the dynamic memory allocation and stack memory usage. Run time allocation of memory heavily depends on the code execution pattern, manually analyzing all possible code execution paths may not be always feasible. So an easier way would be to have an extensive test suit with memory profiling.

Stack for each task in the system can be optimally allocated only by profiling its usage. Here we are not going to detail the various profiling methods, but jot down a strange observation where the stack allocation of a particular function was seen as abnormally high.

Please take a look at the following example code.

Sample Code
Sample Code

As you can see, the above dummy function executes a bunch of switch cases depending on the “code” passed to it, there are some stack variables allocated at the staring of the function (namely x,y & z) . Consider the size of “int” type to be 4 we can comfortably say that the stack allocation at the start of the function will be 4 * 3 = 12 bytes, well instead of this i noticed that the stack allocated by the ARM assembly code is 30. The assembly code looked something like what is given below.

Assembled Code
Assembled Code

Looks like the assembly code is allocating the stack just once after taking into account the maximum stack usage of the function, in other words the first 3 integers (x,y & z) account for 12 and the switch case ‘c’ accounts for the 18 bytes (3 integers and 3 short integers). So the total depth of the stack needs to be at least 30 bytes. Normally we would have expected stack allocation inside each switch case, this means that the assembly code might have had the following sequence included in it.

Non-Optimized Sequence
Non-Optimized Sequence

The above sequence adds more code and more execution overhead, so not really the most optimal way so instead compiler decided that it will allocate the whole stack in one shot. So the execution speed is more and code size also marginally reduced, everything looks good except the fact that now the stack usage seems to have got a hit, how? The stack used inside function “operation_add” would be allocated right after the 30 byte allocation but in the above “non-optimized” sequence we would have had it allocated only after 12 bytes, this means that the optimization in speed comes with a trade-off in the form of stack usage.

This was indeed an interesting observation, the code must have been compiled with optimization setting set for speed instead of memory, hence the compiler conveniently decides not to care much about stack memory. Actually even function inlining will add the stack of the inlined function to that of the caller function, again this is a typical case of speed/memory trade-off, invariably all kinds of  optimization are!

Part 3 — Bluetooth Logical Transports

Content Index:

Introduction : Bluetooth
Part 1: Bluetooth Framework
Part 2: Physical Layer
Part 3: Logical Transports (Link Layer)

 

Physical links are eventually managed by a combination of master – slave relationship and the time slots within the medium. The most cogent way to grasp how these slots are used for transferring data is to explore this within the context of the popular BT use cases given below.

  • Hands free
    &
  • Music Playback (A2DP)
Hands free
Hands free use case is a critical one, this has to be full duplex and real time too which mandates a minimum guaranteed bandwidth otherwise the voice quality will take a hit. For achieving this circuit switched connection like behavior we need to reserve over the air slots for transmitting voice. The following two types of logical transports cater to this very stringent requirement.
  • SCO (Synchronous Connection Oriented) &
  • eSCO (Extended Synchronous Connection Oriented)
When link manager sets up above types of logical links it will make sure that a certain number of slots are reserved for its use, for more clarity please refer to the following figure.
eSCO Window
The above diagram illustrates an eSCO connection with the following attributes.
  • Total six slots are reserved for eSCO.
  • Out of the six slots one slot is for master to slave Tx and one for slave to master.
  • Four slots are reserved for retransmissions
  • There are two unreserved slots between each eSCO window.
During the connection set up phase itself the devices will negotiate the above eSCO connection parameters. There may be multiple unreserved slots between two eSCO windows which may be used for other types of data transfers. eSCO connections and SCO connections work in a similar fashion but in case of SCO there are no re-transmissions and it does not support EDR modulation schemes. Hopefully this has explained how hands free profile is working.
Music Playback (A2DP)
Music play back involves streaming of music files from a device to BT enabled head set. This music data is usually encoded in some form during transfer and later it is decoded and played by the headset, this use case implements a Asynchronous connection oriented transport(ACL). There are no slots reserved for this type of a logical transport, in case of A2DP we can have the headset to be the master or the device to be the master but in either of the case only master can trigger the data transfer. An ACL link is established by exchanging ACL control packets and during the connection establishment the ACL link parameters are negotiated, this will also determine the frequency of the music data transfer, its encoding, the type of packets and this transmission will be frequent enough to guarantee a clear playback.
In case where SCO/eSCO and ACL transports are coexisting between two BT devices, it will make sure that the reserved SCO/eSCO slots are honored and ACL uses only the intermittent unreserved ones.
Now the crux of this logical transports lie in the packet types used, which is given in the table below.
Baseband Packet Types

As you can see above eSCO support more than 1 MBPS rates, this in turn means that EDR (Extended Data Rate) modulation scheme will be used for eSCO but not for SCO, similarly for ACL there are BDR and EDR modulations supported, usage of the packet type will determine the modulation scheme and vice versa is also true.

ACL transport is also used for other profiles like FTP, DUN, SAP etc. One an ACL connection is established to a device then all these application profiles will share this ACL link, multiplexing all these profiles over one link is responsibility of the L2CAP layer.