Released Tuesday, October 19, 2010
Due Wednesday, November 3, 2010, 11:59 PM
In this lab, you will implement a simple disk-based file system. The file system itself will be implemented in micro-kernel fashion, outside the kernel but within its own user-space environment. Other environments access the file system by making IPC requests to this special file system environment.
Use Git to commit your Lab 4 source, fetch the latest version of the course repository, and then create a local branch called lab5 based on our lab5 branch, origin/lab5:
tig% cd ~/CS395T/lab tig% git commit -am 'my solution to lab4' Created commit 734fab7: my solution to lab4 4 files changed, 42 insertions(+), 9 deletions(-) tig% git pull .... tig% git checkout -b lab5 origin/lab5 Branch lab5 set up to track remote branch refs/remotes/origin/lab5. Switched to a new branch "lab5" tig% git merge lab4 Merge made by recursive. kern/env.c | 42 +++++++++++++++++++ 1 files changed, 42 insertions(+), 0 deletions(-) tig%
The main new component for this part of the lab is the file system environment, located in the new fs directory. Scan through all the files in this directory to get a feel for what all is new. Also, there are some new file system-related source files in the user and lib directories, particularly lib/file.c, lib/fd.c, lib/spawn.c, and new global header files inc/fs.h and inc/fd.h. Be sure to scan through all of these files.
You should run the pingpong, primes, and forktree test cases from lab
4 again after merging in the new lab 5 code. You will need to comment
out the ENV_CREATE(fs_fs) line in kern/init.c
because fs/fs.c tries to do some I/O, which JOS does allow
yet. Similarly, temporarily comment out the call to
close_all() in lib/exit.c; this function calls
subroutines that you will
implement later in the lab, and therefore will panic if called. If
your lab 4 code doesn't contain any bugs, the test cases should run
fine. Don't proceed until they work. Don't forget to un-comment
these lines when you start Exercise 1.
If they don't work, use git diff lab4 | more to review all the changes, making sure there isn't any code you wrote for lab4 (or before) missing from lab 5. Make sure that lab 4 still works.
As before, you will need to do all of the regular exercises described in the lab and at least one challenge problem. Additionally, you will need to write up brief answers to the questions posed in the lab and a short (e.g., one or two paragraph) description of what you did to solve your chosen challenge problem. If you implement more than one challenge problem, you only need to describe one of them in the write-up, though of course you are welcome to do more. Place the write-up in a file called answers.txt (plain text) or answers.html (HTML format) in the top level of your lab directory before handing in your work.
When you are ready to hand in your lab code and write-up, create a file called slack.txt noting how many slack hours you have used both for this assignment and in total. (This is to help us agree on the number that you have used.) Then run make turnin in the lab directory. This will first do a make clean to clean out any .o files and executables, and then create a tar file called lab5-handin.tar.gz with the entire contents of your lab directory and submit it via the CS turnin utility. If you submit multiple times, we will take the latest submission and count slack hours accordingly.
The file system that you will work with is much simpler than most "real" file systems, but it is powerful enough to provide the basic features: creating, reading, writing, and deleting files organized in a hierarchical directory structure.
We are (for the moment anyway) developing only a single-user operating system, which provides protection sufficient to catch bugs but not to protect multiple mutually suspicious users from each other. Our file system therefore does not support the UNIX notions of file ownership or permissions. Our file system also currently does not support hard links, symbolic links, time stamps, or special device files like most UNIX file systems do.
Most UNIX file systems divide available disk space into two main types
of regions:
inode regions and data regions.
UNIX file systems assign one inode to each file in the file system;
a file's inode holds critical meta-data about the file
such as its stat attributes and pointers to its data blocks.
The data regions are divided into much larger (typically 8KB or more)
data blocks, within which the file system stores
file data and directory meta-data.
Directory entries contain file names and pointers to inodes;
a file is said to be hard-linked
if multiple directory entries in the file system
refer to that file's inode.
Since our file system will not support hard links,
we do not need this level of indirection
and therefore can make a convenient simplification:
our file system will not use inodes at all
and instead will simply store all of a file's (or sub-directory's) meta-data
within the (one and only) directory entry describing that file.
Both files and directories logically consist of a series of data blocks,
which may be scattered throughout the disk
much like the pages of an environment's virtual address space
can be scattered throughout physical memory.
The file system environment hides the details of block layout,
presenting interfaces for reading and writing sequences of bytes at
arbitrary offsets within files. The file system environment
handles all modifications to directories internally
as a part of performing actions such as file creation and deletion.
Our file system does, however, allow user environments
to read directory meta-data directly
(e.g., with read and write),
which means that user environments can perform directory scanning operations
themselves (e.g., to implement the ls program)
rather than having to rely on additional special calls to the file system.
The disadvantage of this approach to directory scanning,
and the reason that most modern UNIX variants discourage it,
is that it makes application programs dependent
on the format of directory meta-data,
making it difficult to change the file system's internal layout
without changing or at least recompiling application programs as well.
Our file system will use a block size of 4096 bytes, conveniently matching the processor's page size.
File systems typically reserve certain disk blocks at "easy-to-find" locations on the disk (such as the very start or the very end) to hold meta-data describing properties of the file system as a whole, such as the block size, disk size, any meta-data required to find the root directory, the time the file system was last mounted, the time the file system was last checked for errors, and so on. These special blocks are called superblocks.
Our file system will have exactly one superblock,
which will always be at block 1 on the disk.
Its layout is defined by struct Super in inc/fs.h.
Block 0 is typically reserved to hold boot loaders and partition tables,
so file systems generally do not use the very first disk block.
Many "real" file systems maintain multiple superblocks,
replicated throughout several widely-spaced regions of the disk,
so that if one of them is corrupted
or the disk develops a media error in that region,
the other superblocks can still be found and used to access the file system.
In the same way that the kernel must manage the system's physical memory
to ensure that a given physical page is used for only one purpose at a time,
a file system must manage the blocks of storage on a disk
to ensure that a given disk block is used for only one purpose at a time.
In pmap.c you keep the Page structures
for all free physical pages
on a linked list, page_free_list,
to keep track of the free physical pages.
In file systems it is more common to keep track of free disk blocks
using a bitmap rather than a linked list,
because a bitmap is more storage-efficient than a linked list
and easier to keep consistent.
Searching for a free block in a bitmap can take more CPU time
than simply removing the first element of a linked list,
but for file systems this isn't a problem
because the I/O cost of actually accessing the free block after we find it
dominates for performance purposes.
To set up a free block bitmap, we reserve a contiguous region of space on the disk large enough to hold one bit for each disk block. For example, since our file system uses 4096-byte blocks, each bitmap block contains 4096*8=32768 bits, or enough bits to describe 32768 disk blocks. In other words, for every 32768 disk blocks the file system uses, we must reserve one disk block for the block bitmap. A given bit in the bitmap is set if the corresponding block is free, and clear if the corresponding block is in use. The block bitmap in our file system always starts at disk block 2, immediately after the superblock. For simplicity we will reserve enough bitmap blocks to hold one bit for each block in the entire disk, including the blocks containing the superblock and the bitmap itself. We will simply make sure that the bitmap bits corresponding to these special, "reserved" areas of the disk are always clear (marked in-use).
The layout of the meta-data describing a file in our file system is
described by struct File in inc/fs.h. This
meta-data includes the file's name, size, type (regular file or
directory), and pointers to the blocks comprising the file. As
mentioned above, we do not have inodes, so this meta-data is stored in
a directory entry on disk. Unlike in most "real" file systems, for
simplicity we will use this one File structure to
represent file meta-data as it appears
both on disk and in memory.
The f_direct array in struct File contains space
to store the block numbers
of the first 10 (NDIRECT) blocks of the file,
which we call the file's direct blocks.
For small files up to 10*4096 = 40KB in size,
this means that the block numbers of all of the file's blocks
will fit directly within the File structure itself.
For larger files, however, we need a place
to hold the rest of the file's block numbers.
For any file greater than 40KB in size, therefore,
we allocate an additional disk block, called the file's indirect block,
to hold up to 4096/4 = 1024 additional block numbers.
To keep bookkeeping simple, we leave the first 10 numbers in the indirect block
unused. Thus, the 10th block number is the 10th slot in the indirect block
(rather than the 0th, as might be done if we were being very space-efficient).
Our file system therefore allows files to be up to 1024 blocks,
or four megabytes, in size.
To support larger files,
"real" file systems typically support
double- and triple-indirect blocks as well.
A File structure in our file system
can represent either a regular file or a directory;
these two types of "files" are distinguished by the type field
in the File structure.
The file system manages regular files and directory-files
in exactly the same way,
except that it does not interpret the contents of the data blocks
associated with regular files at all,
whereas the file system interprets the contents
of a directory-file as a series of File structures
describing the files and subdirectories within the directory.
The superblock in our file system
contains a File structure
(the root field in struct Super)
that holds the meta-data for the file system's root directory.
The contents of this directory-file
is a sequence of File structures
describing the files and directories located
within the root directory of the file system.
Any subdirectories in the root directory
may in turn contain more File structures
representing sub-subdirectories, and so on.
The goal for this lab is not to have you implement the entire file system, but for you to implement only certain key components. In particular, you will be responsible for reading blocks into the block cache and flushing them back to disk; allocating disk blocks; mapping file offsets to disk blocks; and implementing read, write, and open in the IPC interface. Because you will not be implementing all of the file system yourself, it is very important that you familiarize yourself with the provided code and the various file system interfaces.
Much of the code for the file system has been rewritten for this year, so please let us know if you come across any bugs.
The file system environment in our operating system needs to be able to access the disk, but we have not yet implemented any disk access functionality in our kernel. Instead of taking the conventional "monolithic" operating system strategy of adding an IDE disk driver to the kernel along with the necessary system calls to allow the file system to access it, we will instead implement the IDE disk driver as part of the user-level file system environment. We will still need to modify the kernel slightly, in order to set things up so that the file system environment has the privileges it needs to implement disk access itself.
It is easy to implement disk access in user space this way as long as we rely on polling, "programmed I/O" (PIO)-based disk access and do not use disk interrupts. It is possible to implement interrupt-driven device drivers in user mode as well (the L3 and L4 kernels do this, for example), but it is more difficult since the kernel must field device interrupts and dispatch them to the correct user-mode environment.
The x86 processor uses the IOPL bits in the EFLAGS register to determine whether protected-mode code is allowed to perform special device I/O instructions such as the IN and OUT instructions. Since all of the IDE disk registers we need to access are located in the x86's I/O space rather than being memory-mapped, giving "I/O privilege" to the file system environment is the only thing we need to do in order to allow the file system to access these registers. In effect, the IOPL bits in the EFLAGS register provides the kernel with a simple "all-or-nothing" method of controlling whether user-mode code can access I/O space. In our case, we want the file system environment to be able to access I/O space, but we do not want any other environments to be able to access I/O space at all.
To keep things simple, from now on we will arrange things so that the file system is always user environment 1. (Recall that the idle loop is always user environment 0.)
In the tests that follow, if you fail a test, obj/fs/fs.img is likely to be left inconsistent. Be sure to remove it before running make grade or make qemu.
Exercise 1.
Modify your kernel's environment initialization function,
env_alloc in env.c,
so that it gives environment 1 I/O privilege,
but never gives that privilege to any other environment.
Make sure you can start the file environment without causing a General Protection fault. You should get a 5/100 on make grade.
Do you have to do anything else to ensure that this I/O privilege setting is saved and restored properly when you subsequently switch from one environment to another? Make sure you understand how this environment state is handled.
Read through the files in the new fs directory in the source tree.
The file fs/ide.c implements our minimal PIO-based disk driver.
The file fs/serv.c contains the umain function
for the file system environment.
Note that the GNUmakefile file in this lab sets up QEMU to use the file obj/kern/kernel.img as the image for disk 0 (typically "Drive C" under DOS/Windows) as before, and to use the (new) file obj/fs/fs.img as the image for disk 1 ("Drive D"). In this lab your file system should only ever touch disk 1; disk 0 is used only to boot the kernel. If you manage to corrupt either disk image in some way, you can reset both of them to their original, "pristine" versions simply by typing:
$ rm obj/kern/kernel.img obj/fs/fs.img $ make
or by doing:
$ make clean $ make
Challenge! Implement interrupt-driven IDE disk access, with or without DMA. You can decide whether to move the device driver into the kernel, keep it in user space along with the file system, or even (if you really want to get into the micro-kernel spirit) move it into a separate environment of its own.
In our file system, we will implement a simple "buffer cache" (really just a block cache) with the help of the processor's virtual memory system. The code for the block cache is in fs/bc.c.
Our file system will be limited to handling disks of size 3GB or less.
We reserve a large, fixed 3GB region
of the file system environment's address space,
from 0x10000000 (DISKMAP)
up to 0xD0000000 (DISKMAP+DISKSIZE),
as a "memory mapped" version of the disk.
For example,
disk block 0 is mapped at virtual address 0x10000000,
disk block 1 is mapped at virtual address 0x10001000,
and so on. The diskaddr function in fs/bc.c
implements this translation from disk block numbers to virtual
addresses (along with some sanity checking).
Since our file system environment has its own virtual address space independent of the virtual address spaces of all other environments in the system, and the only thing the file system environment needs to do is to implement file access, it is reasonable to reserve most of the file system environment's address space in this way. It would be awkward for a real file system implementation on a 32-bit machine to do this since modern disks are larger than 3GB. Such a buffer cache management approach may still be reasonable on a machine with a 64-bit address space, such as Intel's Itanium or AMD's Athlon 64 processors.
Of course, it would be unreasonable to read the entire disk into memory, so instead we'll implement a form of demand paging, wherein we only allocate pages in the disk map region and read the corresponding block from the disk in response to a page fault in this region. This way, we can pretend that the entire disk is in memory.
Exercise 2.
Implement the bc_pgfault and flush_block
functions in fs/bc.c.
bc_pgfault is a page fault handler, just like the one
your wrote in the previous lab for copy-on-write fork, except that
its job is to load pages in from the disk in response to a page
fault. When writing this, keep in mind that (1) addr
may not be aligned to a block boundary and (2) ide_read
operates in sectors, not blocks.
The flush_block function should write a block out to disk
if necessary. flush_block shouldn't do anything
if the block isn't even in the block cache (that is, the page isn't
mapped) or if it's not dirty.
We will use the VM hardware to keep track of whether a disk
block has been modified since it was last read from or written to disk.
To see whether a block needs writing,
we can just look to see if the PTE_D "dirty" bit
is set in the vpt entry.
(The PTE_D bit is set by the processor in response to a
write to that page; see 5.2.4.3 in chapter
5 of the 386 reference manual.)
After writing the block to disk, flush_block should clear
the PTE_D bit using sys_page_map.
Use make grade to test your code. Your code should score 20/100.
fs_init function in fs/fs.c is a prime
example of how to use the block cache. After initializing the block
cache, it simply stores pointers into the disk map region in the
super and bitmap global variables. After
this point, we can simply read from these structures as if they were
in memory and our page fault handler will read them from disk as
necessary.
Challenge!
The block cache has no eviction policy. Once a block gets faulted in
to it, it never gets removed and will remain in memory forevermore.
Add eviction to the buffer cache. Using the PTE_A
"accessed" bits in the page tables, you can track approximate usage of
disk blocks without the need to modify every place in the code that
accesses the disk map region. Be careful with dirty blocks.
After fs_init sets the bitmap pointer, we
can treat bitmap as a packed array of bits, one for each
block on the disk. See, for example, block_is_free,
which simply checks whether a given block is marked free in the
bitmap.
Exercise 3.
Use free_block as a model to
implement alloc_block, which should find a free disk
block in the bitmap, mark it used, and return the number of that
block.
When you allocate a block, you should immediately flush
the changed bitmap block to disk with flush_block, to
help file system consistency.
Use make grade to test your code. Your code should score 25/100.
We have provided a variety of functions in fs/fs.c
to implement the basic facilities you will need
to interpret and manage File structures,
scan and manage the entries of directory-files,
and walk the file system from the root
to resolve an absolute pathname.
Read through all of the code in fs/fs.c carefully
and make sure you understand what each function does
before proceeding.
Exercise 4. Implement
file_block_walk
and file_get_block. file_block_walk maps
from a block offset within a file to the pointer for that block in the
struct File or the indirect block, very much like what
pgdir_walk did for page tables.
file_get_block goes one step further and maps to the
actual disk block, allocating a new one if necessary.
Use make grade to test your code. Your code should score 40/100.
file_block_walk and file_get_block are the
workhorses of the file system. For example, file_read
and file_write are little more than the bookkeeping atop
file_get_block necessary to copy bytes between scattered
blocks and a sequential buffer.
Challenge! The file system is likely to be corrupted if it gets interrupted in the middle of an operation (for example, by a crash or a reboot). Implement soft updates or journalling to make the file system crash-resilient and demonstrate some situation where the old file system would get corrupted, but yours doesn't.
The file system server code can be found in fs/serv.c. It
loops in the serve function, endlessly receiving a
request over IPC, dispatching that request to the appropriate handler
function, and sending the result back via IPC. The corresponding
client code can be found in lib/file.c. Here, the
fsipc function handles the common details of sending some
request to the server and receiving the reply. The remaining
functions wrap the available RPC's as regular C functions that handle
bundling up arguments in the request structure, calling
fsipc, and unpacking and returning the results.
Recall that JOS's IPC mechanism lets an environment send a single
32-bit number and, optionally, share a page. To send a request from
the client to the server, we use the 32-bit number for the request
type (the file system server
RPCs are numbered, just as how
syscalls were numbered) and store the arguments to the request in a
union Fsreq on the page shared via the IPC. On the
client side, we always share the page at fsipcbuf; on the
server side, we map the incoming request page at REQVA.
The server also sends the response back via IPC. We use the 32-bit
number for the function's return code. For most RPCs, this is all
they return. FSREQ_READ and FSREQ_STAT also
return data, which they simply write to the page that the client sent
its request on. There's no need to send this page in the response
IPC, since the client shared it with the file system server in the
first place. Also, in its response, FSREQ_OPEN shares with
the client a new "Fd page". We'll return to this RPC shortly.
Exercise 5.
Implement serve_read in fs/serv.c and
file_read in lib/file.c.
serve_read's heavy lifting will be done by
the already-implemented file_read in fs/fs.c (which, in turn, is just a bunch of calls to
file_get_block). serve_read just has to
provide the RPC interface for file reading. Look at the comments and
code in serve_set_size to get a general idea of how the
server functions should be structured.
Likewise, file_read should pack its arguments into
fsipcbuf for serve_read, call
fsipc, and handle the result.
Use make grade to test your code. Your code should score 50/100.
Exercise 6.
Implement serve_write in fs/serv.c and
file_write in lib/file.c.
Use make grade to test your code. Your code should score 60/100.
The functions in lib/file.c are specific to on-disk files,
but UNIX file descriptors are a more general notion that also
encompasses pipes, console I/O, etc. In JOS, each of these device
types has (or will have in later labs) a corresponding struct
Dev, with pointers to the functions that implement
read/write/etc. for that device type. lib/fd.c implements the
general UNIX-like file descriptor interface on top of this. Each
struct Fd indicates its device type, and most of the
functions in lib/fd.c simply dispatch operations to functions
in the appropriate struct Dev.
lib/fd.c also maintains the file descriptor table
region in each application environment's address space, starting at
FDTABLE. This area reserves one page's worth (4KB) of
address space for each of the up to MAXFD (currently 32)
file descriptors that the application can have open at once. At any given
time, a particular file descriptor table page is mapped if and only if
the corresponding file descriptor is in use. Each file descriptor
also has an optional "data page" in the region starting at
FILEDATA, though we won't use that until later labs.
For nearly all interactions with files, user code should go through
the functions in lib/fd.c. The one "public" function in
lib/file.c is open, which constructs a new file
descriptor by opening a named on-disk file.
Exercise 7.
Implement open.
The open function must find an unused file descriptor
using the fd_alloc() function we have provided, make an IPC
request to the file system environment to open the file, and return
the number of the allocated file descriptor. Be sure your code fails
gracefully if the maximum number
of files are already open, or if any of the IPC requests to the file
system environment fails.
Use make grade to test your code. Your code should score 85/100.
Challenge! Change the file system to keep most file meta-data in Unix-style inodes rather than in directory entries, and add support for hard links.
We have given you the code for spawn
which creates a new environment,
loads a program image from the file system into it,
and then starts the child environment running this program.
The parent process then continues running independently of the child.
The spawn function effectively acts like a fork in UNIX
followed by an immediate exec in the child process.
We implemented spawn rather than a
UNIX-style exec because spawn is easier to
implement from user space in "exokernel fashion", without special help
from the kernel. Think about what you would have to do in order to
implement exec in user space, and be sure you understand
why it is harder.
Exercise 8.
spawn relies on the new syscall
sys_env_set_trapframe to initialize the state of the
newly created environment. If you did not already implement
sys_env_set_trapframe in lab 4, then implement it now.
Test your code by running the
user/icode program
from kern/init.c, which will attempt to
spawn /init from the file system.
Use make grade to test your code. Your code should score 100/100.
Challenge!
Implement Unix-style exec.
Challenge!
Implement mmap-style memory-mapped files and
modify spawn to map pages directly from the ELF
image when possible.
Question