by Simon Tatham
Hex editors have been around for a long time, and at the very basic level they are very simple to write. Since they are mostly used for editing files such as executables, which contain a lot of cross-references to particular byte positions in the file, a hex editor need not have an insert mode in order to be useful. And a hex editor without an insert mode is very easy to implement: you simply allocate a large enough array for the input file, and use that as your data structure. The only operation you really need to be able to do efficiently is to jump to a particular byte position, and that's precisely what an array makes easy.
On the other hand, an insert mode can be useful in other circumstances. Not all types of file you might want to edit have the same restrictions as an executable. And as soon as you want your hex editor to have an insert mode, the data structure question becomes much more interesting.
In this article I present an efficient and scalable data structure which supports all the operations needed by a hex editor.
One technique used to support insert mode in editors is to use an array larger than the file size, with a gap in it. The file contents up to the current cursor position are stored at the start of the array; the file contents from the current cursor position to the end are stored at the end of the array; and the gap in the middle moves about as the cursor does.
This makes insertion easy. When the user inserts an extra character, you just add it to one end or other of the gap. On the other hand, moving through the file now becomes a slow operation; it's not noticeable when you're moving by a byte, by a line, or even by a screenful at a time, but as soon as you try to jump to the start or end of the file, or jump to a particular specified file offset, suddenly the editor has to bodily shift enormous amounts of file data from one end of the gap to the other.
Another slightly better option is to use a linked list of small arrays, and to let the arrays vary in size between K and 2K bytes, for some fixed minimum block size K. Inserting a single byte in the middle of a block doesn't cost too much; occasionally the block will grow beyond size 2K and have to be split into two smaller ones, but even that isn't too slow.
Jumping to a particular position, however, is still an O(N) operation using this structure. In practice it isn't too bad, since the length of the linked list is at worst 1/K times the size of the file; but once the file size becomes seriously big, this approach does not scale well.
The common problem in both these methods is that as soon as you make insertion a constant-time operation, seeking to a given byte position becomes linear-time. Whereas in the original array format, of course, seeking was constant-time but insertion became linear-time.
This is where trees come in. Balanced tree structures (any of AVL trees, red-black trees and B-trees) all solve this sort of problem for sorted lists. You can insert an element into a balanced tree in log time, and you can search for a particular element in log time as well. This sounds like the kind of compromise we want: if making insertion constant-time forces seeking to be linear and vice versa, we would prefer to arrange for both to be log-time.
The conventional use of a balanced tree to store a sorted list, however, is not immediately helpful to us. The only criterion we could reasonably sort on would be byte position in the file; and as soon as we store our data as a set of (position, data) pairs, we're back to insertion being linear again, because we would have to alter the position field of every tree element after the insertion point.
Is there anything we can do to our balanced trees to make this work better?
Yes, there is.
Suppose you add an additional field to every node of a balanced tree. In that field, you store a count of the number of elements in or below that node.
Operations which alter the tree (insertion and deletion) now have to make sure these counts remain accurate. This can be done without sacrificing the log-time characteristics of the operations. For example, when you add an element, you increment the count of the node containing it, and then work back up the tree to the root incrementing the counts in all the nodes you go past. Since the height of the tree is O(log N), this only takes you O(log N) time.
So we can add counts to a tree and still maintain it efficiently. What have the counts bought us?
Once we have counts in a tree, they introduce an entirely new way to search the tree. Starting at the root, we can search down the tree by examining the count fields rather than comparing elements as usual; and this allows us to find the Nth item in the tree, for any N, in a single log-time search. For example, suppose the root tree node contains a child with count 54, then an actual element, then a child with count 73. Then:
So now we have a means of finding the Nth item in a tree in a log-time search. This is starting to look promising.
The trouble is, we're still stuck with having some sort of sorting order on the tree. Now we need to deal with that.
The simple answer to the sorting problem is to do away with sorting the tree at all!
Conventional balanced trees have a sorting order because it's used to find elements in the tree, and to know where to add an element. But we don't need a sorting order to find things any more, because we can use a count-based search to jump to the Nth position. Can we also use counts during the tree add operation, to allow us to specify where we want to add our new element?
We can. Tree add algorithms start by searching down the tree to find the position where the new element will be inserted. If we do this search using counts, in exactly the same way described in section 4, then we can add any element we like at any position in the tree. Once we do this, of course, we have to throw out the sorting order completely, and never do another order-based search or insertion again, because they won't work. But that's OK, because we didn't need them anyway.
Now we have exactly what we were after in the first place. We have a data structure which stores an unordered list of items, in such a way that we can insert or delete an item in log time and find the Nth element in log time.
Now we can begin to get more ambitious. One issue we have not addressed yet is cut and paste.
So far I have discussed tree insertion in the assumption that you only ever insert one character at a time into your tree. In fact hex editors need cut and paste just as much as normal text editors do; so we must think about how to insert or remove a larger block of data at a time.
One obvious way is to process each byte individually. A ten-byte cut operation is ten individual deletions, and a ten-byte paste is ten individual insertions. This is fine if you only ever use cut and paste to move tiny chunks of data around a large file, but if you need to move half the file from one place to another, things get more interesting.
The linked-list structure discussed in section 2 would have helped a lot with this problem. Linked lists don't just make it easy to insert or delete one item: they make it just as easy to unlink an enormous chunk of a list once you've found both ends of the chunk, and you can link that chunk in somewhere else easily as well.
It turns out that you can do the same thing with balanced trees. At this point it starts to make a difference what kind of balanced tree you use: all three of AVL, red-black and B-trees support these operations, but the precise methods vary between them. I'm going to use B-trees from here on, because the algorithms are slightly simpler.
What we need are two basic operations. Given a counted, unsorted B-tree containing an unordered list of items, we need to be able to:
This will provide all the operations we need. To unlink a large section from the middle of a tree, we split it in two places and then join the outer two parts back together; to link a large section into the middle of a tree, we split it at the insertion point, join the left half on to the left side of the inserted section, and join the right half on to the right side of the inserted section.
When you add an element to a B-tree, sometimes it ends up increasing the size of a leaf node beyond the size limit. When that happens, you deal with it by splitting the node in two, and transforming the parent node so that where it previously had a single child pointer, it now has two child pointers with an element between them. If that makes the parent node too big as well, you do the same thing again, and so on until you reach the tree root.
Joining two B-trees is therefore reasonably simple, if you have an additional separating element to place in between them. Position the two trees so that their leaf nodes are at the same level; now (usually) one tree will be shorter than the other. So you can add the root of the shorter tree as a sibling of the node next to it in the taller tree; their common parent gains one extra child pointer (pointing at the root of the shorter tree), separated from its neighbour by the additional separating element. If this causes the node to increase beyond the maximum size, just split it in two and propagate up to its parent, just as in the ordinary insertion process.
If the trees were originally the same height, just combine their root nodes into a single larger root node. You need an extra element to go in between the rightmost child pointer of the left-hand root node, and the leftmost child pointer of the right-hand root node; and again, this is where your separating element comes in. Again, if the new root is too big to be a single node, split it in two and create a new root above it.
So it turns out that it's very easy to join two trees together, but the algorithm requires a spare element to go in the middle. However, we normally don't have such a spare element: we just have two trees. This is easily solved, though: we simply start by removing the leftmost element of the right-hand tree using the ordinary tree deletion algorithm. Then we just do the join algorithm, as described above, using the element we just removed as our separator.
To split a B-tree in two: we are given a tree, and a means of searching down the tree to find the split point. (In this application, that will be a numeric position, which we check against the node counts on the way down; in other situations, we might perfectly well want to split an ordinary sorted B-tree in half, so we might have an ordering-based search criterion. It makes no difference.)
We start in the simplest possible way. Start at the root node; decide which of its subtree pointers you are going to descend down; and saw the node in half at that subtree pointer. The two half-nodes thus created will each need a subtree pointer to go on the cut end, but that's OK because we're about to saw the next node down in half as well and they can have half each. So descend to the next node, decide on a split point again, saw that node in half, and put a pointer to each half in the two halves of the parent node.
Once we finish this searching-and-cutting pass, we will have successfully separated our tree into two parts at the required point. However, the result will almost certainly not be a pair of valid B-trees; the chances are that many of the nodes on the cut edges will be below the minimum allowed node size. In fact, if at any point our search criterion made us descend through the endmost subtree pointer in any node, some of those nodes will have no elements in them whatsoever, just a single subtree pointer!
So now we must make a healing pass down the cut edge of each tree, to turn it back into a valid B-tree. We can start by throwing away the root node if it has nothing but a single subtree pointer (which will happen quite often if we split near one end of the original tree, since in that case the output trees will almost certainly need to be of different heights). Keep doing that until we find a real root node.
One child of that node is on the cut edge, so it may be below the minimum size. If it is, we solve this using its (valid) neighbour node. If the neighbour is large, we can move some subtrees over into the undersized node to make two correctly sized nodes; if the neighbour is too small and does not have that many subtrees to spare, we can instead combine the undersized node with its neighbour. (And it turns out you can always do at least one of these: if the neighbour is too large to combine with the undersized node, then it must have enough subtrees for redistribution to give two viable nodes.)
The only interesting case is that combining an undersized node with its neighbour reduces the number of subtrees of their common parent by one. Therefore:
Once we have sorted out each node, we descend to its child on the cut edge, and do the same thing again. Eventually we reach the bottom of the tree and every node is of valid size. Then we do the same thing to the cut edge of the other tree, and we're done.
The splitting and joining algorithms look as if they make cut and paste pretty much trivial. You can split a big chunk out of your editing buffer into a separate cut buffer easily enough; and then you can ‘paste’ it somewhere else by joining it back into the middle of the editing buffer at a different position.
However, in real life, cut and paste isn't that simple. People often want to paste the same data more than once; so you can't just link the cut buffer straight into the editing buffer, because then you don't still have it to link in again next time. You need to copy the cut buffer and link in the copy. Equally, users often want to press Copy rather than Cut, in which case you have to split the buffer tree in two places, copy the middle section, and join all three back together.
Copying a tree, it would seem, is inherently an O(N) operation; there's no way you can copy a tree containing megabytes of data without actually copying all that data.
Or is there?
It turns out that we can do better than this, by adding another annotation field to each tree node. This time, the annotation is a reference count: it counts the number of pointers to the node, either from other tree nodes or from the ‘root’ field in a tree header structure. To begin with, of course, all reference counts are 1.
Reference counts are normally used for garbage collection. In this case, though, I'm going to use them to implement copy-on-write. All of the tree-altering algorithms (insertion and deletion, plus the split and join algorithms described above) will now check the reference count of a node before attempting to modify it. If they find that they need to modify a node with a reference count greater than one, they will not modify it. Instead, they will make a copy of that node, and use the copy in place of the original. The copy links to all the same child nodes as the original, so the reference count in each child must be incremented; and the copied node's parent (or tree header structure) now links to the copy rather than to the original, so the reference count in the original must be decremented. Now we are looking at a node with a reference count of 1, which means nobody else is using it so we can modify it safely.
The effect of this is that it is now a trivial - not merely log-time but constant-time - operation to clone an entire B-tree, no matter how large. We simply create a new tree header structure; we point its root field at the root node of the input tree; and we increment the reference count on that root node.
Once we have cloned a tree like this, we can treat the original and the clone as if they were entirely independent. If you add an element to one of them, for example, then a single string of nodes from the root down to one leaf will be duplicated and modified, but the rest of the trees will still be held in common. You can split either tree into lots of little pieces, or join it into the middle of a larger one, and never affect the data stored in what was once its clone, because every time you touch a node that the other tree is depending on, you make your own copy rather than disturbing it.
This allows us to support really efficient cut and paste in our hex editor. You select a 200Mb chunk and press Copy; the buffer tree is split in two places (in log time), the middle section is cloned (instantly), and the tree is joined back together. You'd hardly know anything was different - but the cut buffer now contains a clone of part of the original buffer, most of which consists of nodes that are still shared with it. And you can paste in as many copies as you like of that chunk, still in no worse than O(log N) time. The best bit is that by the time you've done this a few times and have a file that's 1600Mb longer than it started out, the hex editor isn't actually using up 1600Mb more memory, because most of it is in shared nodes! This technique naturally provides a form of compression as well as being fast.
In all of the above I have been tacitly assuming that the data elements stored in my tree are individual bytes. This would be hideously inefficient if I were using AVL or red-black trees, in which each node contains precisely one element: for every byte of the file being edited, there would be an overhead of two child pointers, a byte count and a reference count. On a normal 32-bit machine, that's 20 bytes per node, not counting overhead from the memory allocator. A factor of twenty is just ridiculous.
B-trees are a bit more flexible, since they can be made to have a large minimum degree. A B-tree with a minimum node size of (say) 512 can contain up to 1023 bytes of data plus 1024 subtree pointers, and those 1023 bytes can be packed together in memory so the overhead is now more like a factor of five. Also, since no node in a B-tree ever changes its height above ground level, you can just not bother to allocate space for the 512 NULL child pointers in your leaf nodes, and since the vast majority of your nodes will be leaf nodes, the structure is now closer to being space-efficient.
There are other improvements one could make. For example, there's no reason why a B-tree really needs to have the same minimum node degree at every level; so you could have low-degree nodes everywhere above the leaf level, and enormous leaf nodes containing 4-8Kb of file data. You could move to B+ trees in which no actual data elements were stored anywhere except in the leaf nodes, thus saving the tiny alignment overheads in the other nodes.
However, there's a better direction to head in. In section 2 I mentioned the idea of using a linked list as the main data structure, and I said that each element of the linked list would be a smallish array of file bytes (between size K and 2K). There's no reason we couldn't do that in our B-tree-based approach: each element stored in the B-tree is no longer a single byte but a small block of bytes. It would mean that our element counts no longer allowed us to jump to the Nth byte, only to the Nth block; but we can fix that by replacing the element count with a byte count, summing the total size of all the blocks in or below a particular tree node. Now, given any byte position, we can do a single log-time search and return a data block plus an offset within that block.
This technique adds work to all operations. Inserting a byte, for example, is now done by finding the block it needs to go into, inserting it in that block, and potentially splitting the block into two and doing an extra tree operation. Splitting and joining buffers involves splitting and joining blocks at each end, and checking to make sure undersized blocks are not created. So what does this technique buy us, that makes it worthwhile over just storing single bytes in the B-tree?
The answer is: once we have a block data structure as our tree element, we can start having different types of block. In particular, we can have a type of block which is a placeholder, containing nothing but a file offset and length. A block of this type indicates ‘at this point in the tree we have N bytes from position P in the original file’. Blocks of this type are exempt from the normal maximum size for normal literal-data blocks.
The effect of this is that we no longer need to read the entire file into memory when we start up. Instead, we just initialise our tree trivially, so that it contains nothing but a single placeholder block, with offset zero and size equal to the initial file size.
Now whenever we need to read data from the tree, and it turns out the data in question is somewhere in a placeholder block, we must refer back to the original input file in order to find the data (and the placeholder block will tell us where in the file to read it from). So before we do any editing, our hex editor is suddenly a low-cost hex file viewer, which just pages back and forth and refers to the disk all the time.
But as soon as we start altering parts of the file, the placeholder block gets broken up into smaller blocks, and literal-data blocks begin to be created in between them. If we cut and paste a section including a placeholder block, then the tree can end up containing placeholder blocks in a strange order; it might (for example) indicate something like ‘the first 192K of the input file; then the literal bytes 5B 49 A7; then 25K of the input file starting from position 12345; then 512K of the input file starting from position 204325’.
Now the hex editor looks as if it's doing exactly the same thing as it did to begin with. I can page around the original file; I can insert, delete, overwrite, cut, copy and paste to my heart's content, and (provided no other process modifies the original file under our feet) the data I am manipulating will remain consistent at all times with the editing operations I have performed. But there wasn't a big delay at startup when the file was loaded in, because most of it wasn't loaded in; and if I list the running processes on my system, the hex editor will not be using memory proportional to the size of the file. It will only be using memory proportional to the changes I've made to the file.
When I save the file, if there are any placeholder blocks remaining in the buffer tree, the hex editor must write out the new version by referring to the original. This is the only remaining operation, apart from searching, that takes time proportional to the size of the file. And there are no remaining operations which take memory proportional to the size of the file.
(There is one thing you need to be careful of. Literal data blocks must be permitted to fall below the minimum size K if there is no literal block next to them to merge with; in particular, this is vital if you are writing a binary file from scratch or you would never be able to give it a size between zero and K. But this raises the possibility that given a pathological sequence of editing operations, your data structure might end up being an interleaving of one-byte literal blocks and one-byte placeholder blocks, giving a huge space overhead. The simplest solution to this is to impose a minimum size of 2K on placeholder blocks, below which you read the relevant piece of file data and convert them into literal blocks; then they can be merged with adjacent blocks and the worst case is no longer terrible.)
We now have a data structure which does pretty much everything you could reasonably ask a hex editor to be able to do, and does it all at a reasonable memory cost and (apart from the two genuinely necessary operations of searching and saving) all in O(log N) time.
The data structure as I have presented it is suitable for use in a high-performance hex editor with an insert mode.
There are a couple more points worth noting.
This structure would need only minor modifications to be an efficient basis for a conventional text editor. In order to do this, you would need to be able to jump quickly to a particular line of the file, which means you'd need a node annotation counting newlines.
In fact, it's possible to do slightly better than that: we can devise a more complex node annotation which tracks the effect of an arbitrary byte sequence on the (line, column) position. Assuming that a physical tab character always advances the cursor to the next multiple of 8 spaces, there are three possibilities:
These three function schemas are closed under composition (i.e. combining any two of them gives another one). Storing one in each node of a buffer tree would provide the ability to search directly to a particular (line, column) position in a single log-time search. So the text editor could treat its buffer as a simple sequence of bytes (or possibly of Unicode characters). This is superior to treating the buffer as a sequence of lines, because it removes the distinction between inserting within a line and inserting data between lines. In particular, cut and paste in a line-based model is fiddly because lines must be spliced together at each end of the pasted region; but cut and paste in this model is as trivial as it was in the hex editor - you just cut a sequence of bytes, paste it somewhere else, and the line/column indexing automatically keeps up no matter what you do.
The only snag is that if you did this, you would probably no longer be able to do the trick with placeholder blocks and lazy file loading; a text editor tends to need to know in advance where all the newlines are in its buffer, so there would probably be no alternative to physically loading the file. But in that, at least, this data structure is no worse than any other.
An undo function in an editor conceptually stores a sequence of previous buffer states, and allows you to return to one of them when you need to.
Usually, this is not actually implemented by storing copies of the entire buffer, since that would be ludicrously wasteful of space! Instead, a journal of changes is kept which allows previous buffer states to be reconstructed by reversing the precise changes made.
One could do that using this data structure, if one wanted to. However, there's another intriguing option. Since cloning an arbitrarily large tree is a cheap operation, you could implement undo by actually storing a sequence of clones of previous buffer states! The cost of this would be nothing like as bad as it would naïvely appear.
It might still not be ideal, though. Every time you clone a tree and the two clones diverge, several nodes must be copied, and if each node contains several blocks of literal data then the cost of maintaining too many buffer clones might still become prohibitive. But it's an interesting possibility regardless.
I've presented a design for a data structure which implements practically every operation required for a hex editor in O(log N) time, apart from one or two which genuinely need to be O(N).
The structure is:
As a result:
Searching must still be linear (there's no alternative to actually reading the data if you need to know anything about its contents), and saving the modified output file is linear (because you actually must physically write out that much data), but everything else can be done in log time.
I've also sketched a means of converting this into a data structure for an ordinary text editor, and suggested interesting implications in the area of undo operations.
Donald Knuth's ‘The Art of Computer Programming’ (Addison-Wesley, ISBN 0201485419) presents at least some of the same ideas as this article. Counted and unsorted trees are mentioned in volume 3; splitting and joining are also described (although Knuth does them on AVL trees, which are significantly more fiddly to split than B-trees; you have to cut the tree into lots of little pieces, and then put them all back together by using the join algorithm repeatedly).
‘Tweak’, a hex editor implementing this data structure, can be downloaded at