Lab 5: File system, Spawn and Shell¶
- Handed out: Tuesday, March 5, 2019
- Due: Thursday, March 14, 2019
In this lab, you will implement
spawn, a library call that loads and
runs on-disk executables. You will then flesh out your kernel and
library operating system enough to run a shell on the console. These
features need a file system, and this lab introduces a simple read/write
Use Git to fetch the latest version of the course repository, and then
create a local branch called
lab5 based on our lab5 branch,
$ cd ~/cs3210/lab $ add git $ git pull Already up-to-date. $ git checkout -b lab5 origin/lab5 Branch lab5 set up to track remote branch refs/remotes/origin/lab5. Switched to a new branch "lab5" $ git merge lab4 Merge made by recursive. ..... $
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
||Code that mainipulates the file system’s on-disk structure.|
||A simple block cache built on top of our user-level page fault handling facility.|
||Minimal PIO-based (non-interrupt-driven) IDE driver code.|
||The file system server that interacts with client environments using file system IPCs.|
||Code that implements the general UNIX-like file descriptor interface.|
||The driver for on-disk file type, implemented as a file system IPC client.|
||The driver for console input/output file type.|
||Code skeleton of the
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
ENV_CREATE(fs_fs) line in
tries to do some I/O, which JOS does not allow yet. Similarly,
temporarily comment out the call to
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
If they don’t work, use git diff lab4 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. Additionally, you will
need to write up brief answers to the questions posed in the lab.
Place the write-up in a file called
in the top level of your
before handing in your work.
File system preliminaries¶
The file system you will work with is much simpler than most “real” file systems including that of xv6 UNIX, 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.
On-Disk File System Structure¶
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
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 allow user
environments to read directory meta-data directly (e.g., with
read), which means that user environments can perform directory
scanning operations themselves (e.g., to implement the
rather than having to rely on additional special calls to the file
system. The disadvantage of this approach to directory scanning, and the
reason 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.
Sectors and Blocks¶
Most disks cannot perform reads and writes at byte granularity and instead perform reads and writes in units of sectors, which today are almost universally 512 bytes each. File systems actually allocate and use disk storage in units of blocks. Be wary of the distinction between the two terms: sector size is a property of the disk hardware, whereas block size is an aspect of the operating system using the disk. A file system’s block size must be a multiple of the sector size of the underlying disk.
The UNIX xv6 file system uses a block size of 512 bytes, the same as the sector size of the underlying disk. Most modern file systems use a larger block size, however, because storage space has gotten much cheaper and it is more efficient to manage storage at larger granularities. 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
The layout of the meta-data describing a file in our
file system is described by
struct File in
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.
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. Our file system therefore
allows files to be up to 1034 blocks, or just over four megabytes, in
size. To support larger files, “real” file systems typically support
double- and triple-indirect blocks as well.
Directories versus Regular Files¶
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
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
representing sub-subdirectories, and so on.
The File System¶
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.
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 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.
i386_init identifies the file system environment by
passing the type
ENV_TYPE_FS to your environment creation function,
env.c, so that it gives the
file system environment 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 pass the “fs i/o” test in make grade.
Question 1. 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? Why?
Note that the
Makefile file in this lab sets up QEMU to use the
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 our
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
$ 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.
The Block Cache¶
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
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+DISKMAX), 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
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.
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.
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.
flush_block function should write a block out to disk if
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
“dirty” bit is set in the
uvpt entry. (The
PTE_D bit is set by
the processor in response to a write to that page; see 184.108.40.206 in
of the 386 reference manual.) After writing the block to disk,
flush_block should clear the
PTE_D bit using
Use make grade to test your code. Your code should pass “check_bc”, “check_super”, and “check_bitmap”.
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 global variable.
After this point, we can simply read from the
super structure as if
they were in memory and our page fault handler will read them from disk
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
“accessed” bits in the page tables, which the hardware sets on any
access to a page, 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.
The Block Bitmap¶
fs_init sets the
bitmap pointer, we can treat
as a packed array of bits, one for each block on the disk. See, for
block_is_free, which simply checks whether a given block is
marked free in the bitmap.
free_block as a model to implement
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
to help file system consistency.
Use make grade to test your code. Your code should now pass “alloc_block”.
We have provided a variety of functions in
fs/fs.c to implement the
basic facilities you will need to interpret and manage
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 and make sure you understand what each
function does before proceeding.
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.
make grade to test your code. Your code should pass “file_open”,
“file_get_block”, and “file_flush/file_truncated/file rewrite”, and
file_get_block are the workhorses of the
file system. For example,
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 interface¶
Now that we have the necessary functionality within the file system environment itself, we must make it accessible to other environments that wish to use the file system. Since other environments can’t directly call functions in the file system environment, we’ll expose access to the file system environment via a remote procedure call, or RPC, abstraction, built atop JOS’s IPC mechanism. Graphically, here’s what a call to the file system server (say, read) looks like
Regular env FS env +---------------+ +---------------+ | read | | file_read | | (lib/fd.c) | | (fs/fs.c) | ...|.......|.......|...|.......^.......|............... | v | | | | RPC mechanism | devfile_read | | serve_read | | (lib/file.c) | | (fs/serv.c) | | | | | ^ | | v | | | | | fsipc | | serve | | (lib/file.c) | | (fs/serv.c) | | | | | ^ | | v | | | | | ipc_send | | ipc_recv | | | | | ^ | +-------|-------+ +-------|-------+ | | +-------------------+
Everything below the dotted line is simply the mechanics of getting a
read request from the regular environment to the file system
environment. Starting at the beginning,
read (which we provide)
works on any file descriptor and simply dispatches to the appropriate
device read function, in this case
devfile_read (we can have more
device types, like pipes).
specifically for on-disk files. This and the other
lib/file.c implement the client side of the FS
operations and all work in roughly the same way, bundling up arguments
in a request structure, calling
fsipc to send the IPC request, and
unpacking and returning the results. The
fsipc function simply
handles the common details of sending a request to the server and
receiving the reply.
The file system server code can be found in
fs/serv.c. It loops in
serve function, endlessly receiving a request over IPC,
dispatching that request to the appropriate handler function, and
sending the result back via IPC. In the read example,
serve_read, which will take care of the IPC details
specific to read requests such as unpacking the request structure and
file_read to actually perform the file read.
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 like how syscalls were numbered)
and store the arguments to the request in a
union Fsipc 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
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
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
the file descriptor page shortly.
serve_read’s heavy lifting will be done by the already-implemented
fs/fs.c (which, in turn, is just a bunch of calls
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.
make grade to test your code. Your code should pass
“serve_open/file_stat/file_close” and “file_read” for a score of
make grade to test your code. Your code should pass “file_write”,
“file_read after file_write”, “open”, and “large file” for a score of
We have given you the code for
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
exec in the child process.
spawn rather than a UNIX-style
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.
spawn relies on the new syscall
sys_env_set_trapframe to initialize the state of the newly created
(don’t forget to dispatch the new system call in
Test your code by running the
user/spawnhello program from
kern/init.c, which will attempt to spawn
/hello from the file
make grade to test your code.
mmap-style memory-mapped files and modify
spawn to map pages directly from the ELF image when possible.
The keyboard interface¶
For the shell to work, we need a way to type at it. QEMU has been
displaying output we write to the CGA display and the serial port, but
so far we’ve only taken input while in the kernel monitor. In QEMU,
input typed in the graphical window appear as input from the keyboard to
JOS, while input typed to the console appear as characters on the serial
kern/console.c already contains the keyboard and serial
drivers that have been used by the kernel monitor since lab 1, but now
you need to attach these to the rest of the system.
kbd_intr to handle trap
serial_intr to handle trap
We implemented the console input/output file type for you, in
serial_intr fill a buffer with
the recently read input while the console file type drains the buffer
(the console file type is used for stdin/stdout by default unless the
user redirects them).
Test your code by running make run-testkbd and type a few lines. The system should echo your lines back to you as you finish them. Try typing in both the console and the graphical window, if you have both available.
Run make run-icode or make run-icode-nox. This will run your kernel and
init, which will set up the
console as file descriptors 0 and 1 (standard input and standard
output). It will then spawn
sh, the shell. You should be able to run
the following commands:
$ echo hello world | cat $ cat lorem |cat $ cat lorem |num $ cat lorem |num |num |num |num |num $ lsfd
Note that the user library routine
cprintf prints straight to the
console, without using the file descriptor code. This is great for
debugging but not great for piping into other programs. To print output
to a particular file descriptor (for example, 1, standard output), use
fprintf(1, "...", ...).
printf("...", ...) is a short-cut for
printing to FD 1. See
user/lsfd.c for examples.
The shell doesn’t support I/O redirection. It would be nice to run
instead of having to type in all the commands in the script by
hand, as you did above. Add I/O redirection for < to
Test your implementation by typing
sh <script into your shell
make run-testshell to test your shell.
testshell simply feeds
the above commands (also found in
fs/testshell.sh) into the shell
and then checks that the output matches
Challenge! Add more features to the shell. Possibilities include (a few require changes to the file system too):
- backgrounding commands (
- multiple commands per line (
ls; echo hi)
- command grouping (
(ls; echo hi) | cat > out)
- environment variable expansion (
- quoting (
echo "a | b")
- command-line history and/or editing
- tab completion
- directories, cd, and a PATH for command-lookup.
- file creation
- ctl-c to kill the running environment
but feel free to do something not on this list.
Your code should pass all tests at this point. As usual, you can grade your submission with make grade and hand it in with make handin.
This completes the lab. As usual, don’t forget to run make grade and
to write up your answers.
Before handing in, use
git status and
git diff to examine your
changes and don’t forget to
git add answers-lab5.txt. When you’re ready,
commit your changes with
git commit -am 'my solutions to lab 5', then
make handin to submit your solution.