If I use inode file system, and I have only 55 free block, than how big is the biggest file, that I can store?
As Inode keep references to all the data blocks where corresponding file data is stored,you can store file which can be accommodated by these 55 free blocks. Also file size depends size of blocks.
Related
Consider a file system that uses contiguous allocation method. For a disk
consists of 100 data blocks, each block is 4KB. What is the maximum and
minimum number of files of size 15KB can it save? Find the number of files the
disk able to support using (1) link allocation and (2) index allocation, assuming
the address is 32 bit
I like to know how data is stored in storage. So what I know about a simple file system organizational struct that contains meta data about a file called inode is that it has two member fields
struct inode {
blkcnt_t i_blocks;
...
loff_t i_size;
}
I am assuming that i_blocks is storing block numbers. but how block numbers are numbered? its of type u64
so the question is if this field contains all the block [numbers] then how they are stored u64 means 64 bit and if I represent each 4 bit relate to block numbers then there are 16 blocks per inode. so for example if i_blocks field is 0b1111 1110.... so 1111 is block number 15 and 1110 is block number 14 and so on. so I like to know if number of bits to represent a block number is 4 bit then there can be only 15 blocks in inode so this way I have block numbers and number of blocks but I still could not field the third field which is >>> what is the base address of data block so for example if inode number is 1111 that correspond to some.txt text file with data hello world then where is the offset of hello world data in storage device. This data offset field array of corresponding inode numbers I could not find. Can any one please direct me to the answer in where I can find the data offset byte in storage medium and it has to be in inode struct?
Short sketch for finding inode number ii:
find the inode block where ii lives: ii/InodesPerBlock; use this as an index into the inodeblocks.
find the index into this block : ii%InodesPerBlock
treat (cast) this location as an Inode, and use the first entry in the blocks[] array as the number of the first data block.
For finding a file, this operation must be precededed by a similar operation for finding the directory entry and finding the file's inodeNumber
NOTE: there are a lot of manifest constants, these can all be found in the superblock
Block size
filesystem Type
size of a blockNumber
number of Inode Blocks (or: number of Inodes)
size of an inode (or: InodesPerBlock)
Number of Data Blocks
Location of the inode containing the root Directory
Location of the freelist containing the unused block numbers
State/backup/...
et cetera ...
NOTE2: this is a simplified scheme. Modern file systems may contain additional structures for efficiency or redundancy. Also: a filesystem may combine more than one physical block into one logical block (e.g. 8*512 -->> 4096)
NOTE3: the superblock is located at blocknumber=0 (at least in UFS) That means that 0 can be used as sentinel value for blocknumbers referring to actual (non-root) blocks. So, the blocknumber arrays inside the inodes can be initialized to all-zeros.
NOTE4: these structures all reside on disk. The kernel may maintain (it will!) additional, similar structures in memory. These structures will refer to both themselves (using pointers, or offsets, or indices) or they will refer to disk blocks (numbers).
I'm learning about in-kernel data transferring between two file descriptors in Linux and came across something I cannot understand. Here is the quote from copy_file_range manpage
copy_file_range() gives filesystems an opportunity to implement "copy
acceleration" techniques, such as the use of reflinks (i.e., two or
more i-nodes that share pointers to the same copy-on-write disk
blocks) or server-side-copy
I used to think of index nodes as something that is returned by stat/statx syscall. The st_ino type is typedefed here as
typedef unsigned long __kernel_ulong_t;
So what does it ever mean "two or more i-nodes that share pointers to the same copy-on-write disk blocks"?
According to my understanding the fact that copy_file_range do not need to pass the data through the user-mode means the kernel doesn't have to load the data from the disk at all (it still might but it doesn't have to) and this allows further optimization by pushing the operation down the file-system stack. This covers the case of the server-side-copy over NFS.
The actual answers about the other optimization starts with an intro into how files are stored, you may skip it if you already know that.
There are 3 layers in how files are stored in a typical Linux FS:
The file entry in some directory (which is itself a file containing a list of such entries). Such entry essentially maps file name to some inode. It is done by storing the inode-number aka st_ino which is effectively a pointer to the inode in some table.
The inode that contains some shared (see further) metadata (as the one returned by stat) and some pointer(s) to data block(s) that store the actual file contents.
The actual data blocks
So for example a hard-link is a record in some directory that points to the same inode as the "original" file (and incrementing the "link counter" inside the inode). This means that only file names (and possibly directories) are different, all the rest of the data and meta-data is shared between hard-links. Note that creating a hard link is a very fast way to copy a file. The only drawback is that both files now are bound to share their contents forever so this is not a true copy. But if we used some copy-on-write method to fix the "write" part, it would work very nice. This is what some FSes (such as Btrfs) support via reflinks.
The idea of this copy-on-wrote trick is that you can create a new inode with new appropriate metadata but still share the same data blocks. You also add cross-references between the two inodes in the "invisible" part of the inode metadata so they know they share the data blocks. Obviously this operation is very fast comparing to the real copying. And again as long as the files are only read, everything works perfectly. But unlike hard-link we can deal with writes treating them as independent as well. When some write is performed, the FS checks if the file (or rather the inode) is really the only owner of the data blocks and else copies the data before writing to it. Depending on the FS implementation it can copy the whole file on the first write or it can store some more detailed metadata and only copy the blocks that have to be modified and still share the rest between the files. In the later case blocks might not need to be copied at all if the write size is more than a block.
So the simplest trick copy_file_range() can do is to check if the whole file is actually being copied and if so, to perform the reflink trick described above (obviously if the FS supports it).
Some more advanced optimizations are also possible if the FS supports more detailed meta-data on data blocks. Assume you copy first N bytes from the start of the file into a new file. Then the FS can just share the starting blocks and probably has to copy only the last one that is not fully copied.
Using the FatFs and their API, I am trying to pre-allocate file system space for the remainder of the drive to write files of unknown sizes. At the end of the file write process, any unused space is then truncated (f_truncate). However, I'm having a problem allocating enough space when the file system become fragmented after deletion of some files. To overcome this problem, I would like to allocate only enough space as the largest contiguous block of memory on the file system.
In the FatFS API, there are functions to get the amount of free space left on the device, which is f_getfree. There is also the f_expand function, which takes in a size in bytes and returns whether or not a free contiguous block of memory exists for that given size.
Is there an efficient way to calculate the largest free contiguous block of memory available? I'm trying to avoid any sort of brute force "guess and check" method. Thanks
One way would be to create your own extension to the API to count contiguous FAT entries across all the sectors.
Without modifying the API, you could use f_lseek(). Write open a file, use f_lseek() to expand the size of the file until 'disk full'(end of contiguous space). This would need to be repeated with new files until all of the disk was allocated. Pick from this the maximum allocated file, and delete the others.
I need a function which will create a file with fixed size in linux. Function like truncate or fopen,fseek,fclose, is not a solution because, they will fill opened file with zeros, but it is not necessary and I have no time for this. So is there some function, which will only open a file with fixed length and not fill buffer?
Thanks in advance.
The system call truncate(2) doesn't fill the file with zeros. It simply advances the file's reported size and leaves holes in it.
When you read from it, you do get zeros, but that's just a convenience of the OS.
The truncate() and ftruncate() functions cause the regular file named
by path or referenced by fd to be truncated to a size of precisely
length bytes.
If the file previously was shorter, it is extended, and the extended
part reads as null bytes ('\0').
About holes (from TLPI):
The existence of holes means that a file’s nominal size may be larger
than the amount of disk storage it utilizes (in some cases,
considerably larger).
Filesystems and holes:
Rather than allocate blocks of null bytes for the holes in a file, the
file system can just mark (with the value 0) appropriate pointers in
the i-node and in the indirect pointer blocks to indicate that they
don't refer to actual disk blocks.
As Per Johansson notes, this is dependent of the filesystem.
Most native UNIX file systems support the concept of file holes, but
many nonnative file systems (e.g., Microsoft’s VFAT) do not. On a
file system that doesn’t support holes, explicit null bytes are
written to the file.