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Text Processing, Part 2: Oh, Inverted Index
This is the second part of my text processing series. In this blog, we'll look into how text documents can be stored in a form that can be easily retrieved by a query. I'll used the popular open source Apache Lucene index for illustration. There are two main processing flow in the system ... Document indexing: Given a document, add it into the index Document retrieval: Given a query, retrieve the most relevant documents from the index. The following diagram illustrate how this is done in Lucene. Index Structure Both documents and query is represented as a bag of words. In Apache Lucene, "Document" is the basic unit for storage and retrieval. A "Document" contains multiple "Fields" (also call zones). Each "Field" contains multiple "Terms" (equivalent to words). To control how the document will be indexed across its containing fields, a Field can be declared in multiple ways to specified whether it should be analyzed (a pre-processing step during index), indexed (participate in the index) or stored (in case it needs to be returned in query result). Keyword (Not analyzed, Indexed, Stored) Unindexed (Not analyzed, Not indexed, Stored) Unstored (Analyzed, Indexed, Not stored) Text (Analyzed, Indexed, Stored) The inverted index is a core data structure of the storage. It is organized as an inverted manner from terms to the list of documents (which contain the term). The list (known as posting list) is ordered by a global ordering (typically by document id). To enable faster retrieval, the list is not just a single list but a hierarchy of skip lists. For simplicity, we ignore the skip list in subsequent discussion. This data structure is illustration below based on Lucene's implementation. It is stored on disk as segment files which will be brought to memory during the processing. The above diagram only shows the inverted index. The whole index contain an additional forward index as follows. Document indexing Document in its raw form is extracted from a data adaptor. (this can be making an Web API to retrieve some text output, or crawl a web page, or receiving an HTTP document upload). This can be done in a batch or online manner. When the index processing start, it parses each raw document and analyze its text content. The typical steps includes ... Tokenize the document (breakdown into words) Lowercase each word (to make it non-case-sensitive, but need to be careful with names or abbreviations) Remove stop words (take out high frequency words like "the", "a", but need to careful with phrases) Stemming (normalize different form of the same word, e.g. reduce "run", "running", "ran" into "run") Synonym handling. This can be done in two ways. Either expand the term to include its synonyms (ie: if the term is "huge", add "gigantic" and "big"), or reduce the term to a normalized synonym (ie: if the term is "gigantic" or "huge", change it to "big") At this point, the document is composed with multiple terms. doc = [term1, term2 ...]. Optionally, terms can be further combined into n-grams. After that we count the term frequency of this document. For example, in a bi-gram expansion, the document will become ... doc1 -> {term1: 5, term2: 8, term3: 4, term1_2: 3, term2_3:1} We may also compute a "static score" based on some measure of quality of the document. After that, we insert the document into the posting list (if it exist, otherwise create a new posting list) for each terms (all n-grams), this will create the inverted list structure as shown in previous diagram. There is a boost factor that can be set to the document or field. The boosting factor effectively multiply the term frequency which effectively affecting the importance of the document or field. Document can be added to the index in one of the following ways; inserted, modified and deleted. Typically the document will first added to the memory buffer, which is organized as an inverted index in RAM. When this is a document insertion, it goes through the normal indexing process (as I described above) to analyze the document and build an inverted list in RAM. When this is a document deletion (the client request only contains the doc id), it fetches the forward index to extract the document content, then goes through the normal indexing process to analyze the document and build the inverted list. But in this case the doc object in the inverted list is labeled as "deleted". When this is a document update (the client request contains the modified document), it is handled as a deletion followed by an insertion, which means the system first fetch the old document from the forward index to build an inverted list with nodes marked "deleted", and then build a new inverted list from the modified document. (e.g. If doc1 = "A B" is update to "A C", then the posting list will be {A:doc1(deleted) -> doc1, B:doc1(deleted), C:doc1}. After collapsing A, the posting list will be {A:doc1, B:doc1(deleted), C:doc1} As more and more document are inserted into the memory buffer, it will become full and will be flushed to a segment file on disk. In the background, when M segments files have been accumulated, Lucene merges them into bigger segment files. Notice that the size of segment files at each level is exponentially increased (M, M^2, M^3). This maintains the number of segment files that need to be search per query to be at the O(logN) complexity where N is the number of documents in the index. Lucene also provide an explicit "optimize" call that merges all the segment files into one. Here lets detail a bit on the merging process, since the posting list is already vertically ordered by terms and horizontally ordered by doc id, merging two segment files S1, S2 is basically as follows Walk the posting list from both S1 and S2 together in sorted term order. For those non-common terms (term that appears in one of S1 or S2 but not both), write out the posting list to a new segment S3. Until we find a common term T, we merge the corresponding posting list from these 2 segments. Since both list are sorted by doc id, we just walk down both posting list to write out the doc object to a new posting list. When both posting lists have the same doc (which is the case when the document is updated or deleted), we pick the latest doc based on time order. Finally, the doc frequency of each posting list (of the corresponding term) will be computed. Document retrieval Consider a document is a vector (each term as the separated dimension and the corresponding value is the tf-idf value) and the query is also a vector. The document retrieval problem can be defined as finding the top-k most similar document that match a query, where similarity is defined as the dot-product or cosine distance between the document vector and the query vector. tf-idf is a normalized frequency. TF (term frequency) represents how many time the term appears in the document (usually a compression function such as square root or logarithm is applied). IDF is the inverse of document frequency which is used to discount the significance if that term appears in many other documents. There are many variants of TF-IDF but generally it reflects the strength of association of the document (or query) with each term. Given a query Q containing terms [t1, t2], here is how we fetch the corresponding documents. A common approach is the "document at a time approach" where we traverse the posting list of t1, t2 concurrently (as opposed to the "term at a time" approach where we traverse the whole posting list of t1 before we start the posting list of t2). The traversal process is described as follows ... For each term t1, t2 in query, we identify all the corresponding posting lists. We walk each posting list concurrently to return a sequence of documents (ordered by doc id). Notice that each return document contains at least one term but can also also contain multiple terms. We compute the dynamic score which is dot product of the query to document vector. Notice that we typically don't concern the TF/IDF of the query (which is short and we don't care the frequency of each term). Therefore we can just compute the sum up all the TF score of the posting list that has a match term after dividing the IDF score (at the head of each posting list). Lucene also support query level boosting where a boost factor can be attached to the query terms. The boost factor will multiply the term frequency correspondingly. We also look up the static score which is purely based on the document (but not the query). The total score is a linear combination of static and dynamic score. Although the score we used in above calculation is based on computing the cosine distance between the query and document, we are not restricted to that. We can plug in any similarity function that make sense to the domain. (e.g. we can use machine learning to train a model to score the similarity between a query and a document). After we compute a total score, we insert the document into a heap data structure where the topK scored document is maintained. Here the whole posting list will be traversed. In case of the posting list is very long, the response time latency will be long. Is there a way that we don't have to traverse the whole list and still be able to find the approximate top K documents ? There are a couple strategies we can consider. Static Score Posting Order: Notice that the posting list is sorted based on a global order, this global ordering provide a monotonic increasing document id during the traversal that is important to support the "document at a time" traversal because it is impossible to visit the same document again. This global ordering, however, can be quite arbitrary and doesn't have to be the document id. So we can pick the order to be based on the static score (e.g. quality indicator of the document) which is global. The idea is that we traverse the posting list in decreasing magnitude of static score, so we are more likely to visit the document with the higher total score (static + dynamic score). Cut frequent terms: We do not traverse the posting list whose term has a low IDF value (ie: the term appears in many documents and therefore the posting list tends to be long). This way we avoid to traverse the long posting list. TopR list: For each posting list, we create an extra posting list which contains the top R documents who has the highest TF (term frequency) in the original list. When we perform the search, we perform our search in this topR list instead of the original posting list. Since we have multiple inverted index (in memory buffer as well as the segment files at different levels), we need to combine the result them. If termX appears in both segmentA and segmentB, then the fresher version will be picked. The fresher version is determine as follows; the segment with a lower level (smaller size) will be considered more fresh. If the two segment files are at the same level, then the one with a higher number is more fresh. On the other hand, the IDF value will be the sum of the corresponding IDF of each posting list in the segment file (the value will be slightly off if the same document has been updated, but such discrepancy is negligible). However, the processing of consolidating multiple segment files incur processing overhead in document retrieval. Lucene provide an explicit "optimize" call to merge all segment files into one single file so there is no need to look at multiple segment files during document retrieval. Distributed Index For large corpus (like the web documents), the index is typically distributed across multiple machines. There are two models of distribution: Term partitioning and Document partitioning. In document partitioning, documents are randomly spread across different partitions where the index is built. In term partitioning, the terms are spread across different partitions. We'll discuss document partitioning as it is more commonly used. Distributed index is provider by other technologies that is built on Lucene, such as ElasticSearch. A typical setting is as follows ... In this setting, machines are organized as columns and rows. Each column represent a partition of documents while each row represent a replica of the whole corpus. During the document indexing, first a row of the machines is randomly selected and will be allocated for building the index. When a new document crawled, a column machine from the selected row is randomly picked to host the document. The document will be sent to this machine where the index is build. The updated index will be later propagated to the other rows of replicas. During the document retrieval, first a row of replica machines is selected. The client query will then be broadcast to every column machine of the selected row. Each machine will perform the search in its local index and return the TopM elements to the query processor which will consolidate the results before sending back to client. Notice that K/P < M < K, where K is the TopK documents the client expects and P is the number of columns of machines. Notice that M is a parameter that need to be tuned. One caveat of this distributed index is that as the posting list is split horizontally across partitions, we lost the global view of the IDF value without which the machine is unable to calculate the TF-IDF score. There are two ways to mitigate that ... Do nothing: here we assume the document are evenly spread across different partitions so the local IDF represents a good ratio of the actual IDF. Extra round trip: In the first round, query is broadcasted to every column which returns its local IDF. The query processor will collected all IDF response and compute the sum of the IDF. In the second round, it broadcast the query along with the IDF sum to each column of machines, which will compute the local score based on the IDF sum.
February 26, 2013
by Ricky Ho
· 9,611 Views
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Solving RPM installation conflicts
This post comes from Ignacio Nin at the MySQL Performance Blog. Lately we’ve had many reports of the RPM packages for CentOS 5 (mostly) and CentOS 6 having issues when installing different combinations of our products, particularly with Percona Toolkit. Examples of bugs related to these issues are lp:1031427 and lp:1051874. These problems arise when trying to install a package from the distribution that is linked against the version of libmysqlclient.so shipped by the distribution (libmysqlclient.so.15 for CentOS 5/libmysqlclient.so.16 for CentOS 6) and a version of Percona Server that depends on another version of libmysqlclient.so, usually more recent. Bug lp:1031427 is an example of this, and shows how the packages would conflict when trying to install libmysqlclient.so. For example, when installing php-mysql alongside PS 5.5 in CentOS 6: # yum -q install Percona-Server-server-55 php-mysql Installing: Percona-Server-server-55 x86_64 5.5.29-rel29.4.401.rhel6 percona 15 M php-mysql x86_64 5.3.3-14.el6_3 updates 79 k Installing for dependencies: Percona-Server-client-55 x86_64 5.5.29-rel29.4.401.rhel6 percona 7.0 M Percona-Server-shared-51 x86_64 5.1.67-rel14.3.506.rhel6 percona 2.8 M Percona-Server-shared-55 x86_64 5.5.29-rel29.4.401.rhel6 percona 787 k Transaction Summary ===================================================================================================================================================== Install 5 Package(s) Is this ok [y/N]: y Transaction Check Error: file /usr/lib64/libmysqlclient.so conflicts between attempted installs of Percona-Server-shared-51-5.1.67-rel14.3.506.rhel6.x86_64 and Percona-Server-shared-55-5.5.29-rel29.4.401.rhel6.x86_64 file /usr/lib64/libmysqlclient_r.so conflicts between attempted installs of Percona-Server-shared-51-5.1.67-rel14.3.506.rhel6.x86_64 and Percona-Server-shared-55-5.5.29-rel29.4.401.rhel6.x86_64 The traditional solution for this situation was to provide a special package, Percona-Server-shared-compat (modeled after upstream’s MySQL-shared-compat) which would contain ALL versions of libmysqlclient.so.* together and wouldn’t conflict. Probably some of you are familiar with this approach. # yum -q install Percona-Server-server-55 Percona-Server-shared-compat php-mysql Installing: Percona-Server-server-55 x86_64 5.5.29-rel29.4.401.rhel6 percona 15 M Percona-Server-shared-compat x86_64 5.5.29-rel29.4.401.rhel6 percona 3.4 M php-mysql x86_64 5.3.3-14.el6_3 updates 79 k Installing for dependencies: Percona-Server-client-55 x86_64 5.5.29-rel29.4.401.rhel6 percona 7.0 M Percona-Server-shared-55 x86_64 5.5.29-rel29.4.401.rhel6 percona 787 k Transaction Summary ===================================================================================================================================================== Install 5 Package(s) Notice how PS-shared-compat installs along the -shared package, providing the older libmysqlclient.so.16 required by php-mysql. However, this has proved non-intuitive and problematic, since the shared-compat package wouldn’t get selected unless explicitely installed — and many of our users would rather have it “just work” without requiring additional knowledge of what the particular workaround was, etc.. We’re now trying a solution in which our -shared packages won’t conflict anymore at libmysqlclient.so, so we are able to install them side-by-side, modelled after the mysql-libs packages provided by CentOS/Redhat. So even if the user wants to install PS 5.5 alongside packages that depend on 5.1/5.0, the -shared packages will work together. For example installing 5.5 and postfix in CentOS: # yum -q install Percona-Server-server-55 postfix Installing: Percona-Server-server-55 x86_64 5.5.29-rel29.4.402.rhel5 percona-testing 19 M postfix x86_64 2:2.3.3-6.el5 base 3.8 M Installing for dependencies: Percona-SQL-shared-50 x86_64 5.0.92-b23.89.rhel5 percona-testing 1.8 M Percona-Server-client-55 x86_64 5.5.29-rel29.4.402.rhel5 percona-testing 9.1 M Percona-Server-shared-55 x86_64 5.5.29-rel29.4.402.rhel5 percona-testing 993 k … and this will install without problems. Additionally, this has the advantage of allowing an upgrade from 5.1 to 5.5 without uninstalling any software that depended on the old version. # rpm -qa | grep ^Percona Percona-Server-client-51-5.1.67-rel14.3.507.rhel6.x86_64 Percona-Server-shared-51-5.1.67-rel14.3.507.rhel6.x86_64 Percona-Server-server-51-5.1.67-rel14.3.507.rhel6.x86_64 In this case only Percona-Server-client-51 and Percona-Server-server-51 need be removed, allowing any package that depends on Percona-Server-shared-51 (providing libmysqlclient.so.16) to remain installed. After the server and client packages are uninstalled, you can install PS 5.5 without conflict. The current package candidates for versions 5.0.92 (which required an update), 5.1.67-14.3 and 5.5.29-29.4 can be tested from the percona-testing repository. We encourage you to try these out and send us your feedback and/or file any bugs you find. Installation instructions for Percona Testing repositories. We’re aiming to include these fixes in our next releases of 5.1 and 5.5. Percona Toolkit users in particular will enjoy this update since it’ll mean no more trouble when installing it from repository!
February 25, 2013
by Peter Zaitsev
· 7,954 Views
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Building SOLID Databases: Dependency Inversion and Robust DB Interfaces
Dependency inversion is the idea that interfaces should depend on abstractions not on specifics. According to Wikipedia, the principle states: A. High-level modules should not depend on low-level modules. Both should depend on abstractions. B. Abstractions should not depend upon details. Details should depend upon abstractions. Of course the second part of this principle is impossible if read literally. You can't have an abstraction until you know what details are to be covered, and so the abstraction and details are both co-dependent. If the covered details change sufficiently the abstraction will become either leaky or inadequate and so it is worth seeing these as intertwined to some extent. The focus on abstraction is helpful because it suggests that the interface contract should be designed in such a way that neither side really has to understand any internal details of the other in order to make things work. Both sides depend on well-encapsulated API's and neither side has to worry about what the other side is really doing. This is what is meant by details depending on abstractions rather than the other way around. This concept is quite applicable beyond object oriented programming because it covers a very basic aspect of API contract design, namely how well an API should encapsulate behavior. This principle is first formulated in its current form in the object oriented programming paradigm but is generally applicable elsewhere. SQL as an Abstraction Layer, or Why RDBMS are Still King There are plenty of reasons to dislike SQL, such as the fact that nulls are semantically ambiguous. As a basic disclaimer I am not holding SQL up to be a paragon of programming languages or even db interfaces, but I think it is important to discuss what SQL does right in this regard. SQL is generally understood to be a declarative language which approximates relational mathematics for database access purposes. With SQL, you specify what you want returned, not how to get it, and the planner determines the best way to get it. SQL is thus an interface language rather than a programming language per se. With SQL, you can worry about the logical structure, leaving the implementation details to the db engine. SQL queries are basically very high level specifications of operations, not detailed descriptions of how to do something efficiently. Even update and insert statements (which are by nature more imperative than select statements) leave the underlying implementation entirely to the database management system. I think that this, along with many concessions the language has made to real-world requirements (such as bags instead of sets and the addition of ordering to bags) largely account for the success of this language. SQL, in essence, encapsulates a database behind a mature mathematical, declarative model in the same way that JSON and REST do (in a much less comprehensive way) in many NoSQL db's. In essence SQL provides encapsulation, interface, and abstraction in a very full-featured way and this is why it has been so successful. SQL Abstraction as Imperfect One obvious problem with treating SQL as an abstraction layer in its own right is that one is frequently unable to write details in a way that is clearly separate from the interface. Often storage tables are hit directly, and therefore there is little separation between logical detail and logical interface, and so this can break down when database complexity reaches a certain size. Approaches to managing this problem include using stored procedures or user defined functions, and using views to encapsulate storage tables. Stored Procedures and User Defined Functions Done Wrong Of the above methods, stored procedures and functional interfaces have bad reputations frequently because of bad experiences that many people have with them. These include developers pushing too much logic into stored procedures, and the fact that defining functional interfaces in this way usually produces a very tight binding between database code and application code, often leading to maintainability problems. The first case is quite obvious, and includes the all-too-frequent case of trying to send emails directly from stored procedures (always a bad idea). This mistake leads to certain types of problems, including the fact that ACID-compliant operations may be mixed with non-ACID-compliant ones, leading to cases where a transaction can only be partially rolled back. Oops, we didn't actually record the order as shipped, but we told the customer it was..... MySQL users will also note this is an argument against mixing transactional and nontransactional backend table types in the same db..... However that problem is outside the scope of this post. Additionally, MySQL is not well suited for many applications against a single set of db relations. The second problem, though, is more insidious. The traditional way stored procedures and user defined functions are typically used, the application has to be deeply aware of the interface to the database, but the rollout for these aspects is different leading to the possibility or service interruptions, and a need to very carefully and closely time rollout of db changes with application changes. As more applications use the database, this becomes harder and the chance of something being overlooked becomes greater. For this reason the idea that all operations must go through a set of stored procedures is a decision fraught with hazard as the database and application environment evolves. Typically it is easier to manage backwards-compatibility in schemas than it is in functions and so a key question is how many opportunities you have to create new bugs when a new column is added. There are, of course, more hazards which I have dealt with before, but the point is that stored procedures are potentially harmful and a major part of the reason is that they usually form a fairly brittle contract with the application layer. In a traditional stored procedure, adding a column to be stored will require changing the number of variables in the stored procedure's argument list, the queries to access it, and each application's call to that stored procedure. In this way, they provide (in the absence of other help) at best a leaky abstraction layer around the database details. This is the sort of problem that dependency inversion helps to avoid. Stored Procedures and User Defined Functions Done Right Not all stored procedures are done wrong. In the LedgerSMB project we have at least partially solved the abstraction/brittleness issue by looking to web services for inspiration. Our approach provides an additional mapping layer and dynamic query generation around a stored procedure interface. By using a service locator pattern, and overloading the system tables in PostgreSQL as the service registry, we solve the problem of brittleness. Our approach of course is not perfect and it is not the only possibility. One shortcoming is that our approach is that the invocation of the service locator is relatively spartan. We intend to allow more options there in the future. However one thing I have noticed is the fact that there are far fewer places where bugs can hide and therefore faster and more robust development takes place. Additionally a focus on clarity of code in stored procedures has eliminated a number of important performance bottlenecks, and it limits the number of places where a given change propagates to. Other Important Options in PostgreSQL Stored procedures are not the only abstraction mechanisms available from PostgreSQL. In addition to views, there are also other interesting ways of using functions to accomplish this without insisting that all access goes through stored procedures. In addition these methods can be freely mixed to produce very powerful, intelligent database systems. Such options include custom types, written in C, along with custom operators, functions and the like. These would then be stored in columns and SQL can be used to provide an abstraction layer around the types. In this way SQL becomes the abstraction and the C programs become the details. A future post will cover the use of ip4r in network management with PostgreSQL db's as an example of what can be done here. Additionally, things like triggers and notifications can be used to ensure that appropriate changes trigger other changes in the same transaction or, upon transaction commit, hand off control to other programs in subsequent transactions (allowing for independent processing and error control for things like sending emails). Recommendations Rather than specific recommendations, the overall point here is to look at the database itself as a an application running in an application server (the RDBMS) and design it as an application with an appropriate API. There are many ways to do this, from writing components in C and using SQL as an abstraction mechanism to writing things in SQL and using stored procedures as a mechanism. One could even write code in SQL and still use SQL as an abstraction mechanism. The key point however is to be aware of the need for discoverable abstraction, a need which to date things like ORMs and stored procedures often fill very imperfectly. A well designed db with appropriate abstraction in interfaces, should be able to be seen as an application in its own right, engineered as such, and capable of serving multiple client apps through a robust and discoverable API. As with all things, it starts by recognizing the problems and putting solutions as priorities from the design stage onward.
February 19, 2013
by Chris Travers
· 5,368 Views
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Neo4j/Cypher: SQL Style GROUP BY Functionality
As I mentioned in a previous post I’ve been playing around with some football related data over the last few days and one query I ran (using cypher) was to find all the players who’ve been sent off this season in the Premiership. The model in the graph around sending offs looks like this: My initial query looked like this: START player = node:players('name:*') MATCH player-[:sent_off_in]-game-[:in_month]-month RETURN player.name, month.name First we get the names of all the players which are stored in an index and then we follow relationships to the games they were sent off in and then find which months those games were played in. That query returns: +----------------------------+ | player.name | month.name | +----------------------------+ | "Jenkinson" | "February" | | "Chico" | "September" | | "Odemwingie" | "September" | | "Agger" | "August" | | "Cole" | "December" | | "Whitehead" | "August" | ... +----------------------------+ I thought it’d be interesting to see how many sending offs there were in each month which we’d achieve in SQL by making use of a GROUP BY. cypher has a bunch of aggregation functions which allow us to achieve the same outcome. In our case we want to use the COUNT function and we want our grouping key to be the month of the year so we need to include that as part of our RETURN statement as well: START player = node:players('name:*') MATCH player-[:sent_off_in]-game-[:in_month]-month RETURN COUNT(player.name) AS numberOfReds, month.name ORDER BY numberOfReds DESC which returns: +----------------------------+ | numberOfReds | month.name | +----------------------------+ | 7 | "October" | | 6 | "December" | | 4 | "September" | | 4 | "November" | | 3 | "August" | | 2 | "January" | | 2 | "February" | +----------------------------+ As far as I can tell anything which isn’t an aggregate function is used as part of the grouping key which means we could include more than one field in our grouping key. This isn’t particularly relevant for us for this particular query but would become useful if we add the teams that the players play for. I extended the graph to included a player’s statistics for each game which also includes a relationship indicating which team they played for in a specific game. The model now looks like this: It does now look quite a bit more complicated but this was the best way I could think of modelling player specific details for a match. I couldn’t see another way of modelling the fact that a player played for a certain team in a match which I want to use for some other queries but if you can see a simpler way please let me know. To get a list of the red cards and the name of the team the offender played for we can write the following query: START player = node:players('name:*') MATCH player-[:sent_off_in]-game-[:in_month]-month, game-[:in_match]-stats-[:stats]-player, stats-[:played_for]-team RETURN player.name, month.name, team.name ORDER BY month.name The original query traversed a path from a player to games they were sent off in and then from the games to the month the game was played in. We’ve now added a traversal from the game to the game stats for that player and we also traverse from the game stats to the team node that the player played for in that game. When we run this we get the following results: +--------------------------------------------+ | player.name | month.name | team.name | +--------------------------------------------+ | "Agger" | "August" | "Liverpool" | | "Whitehead" | "August" | "Stoke" | ... | "Shotton" | "December" | "Stoke" | | "Nzonzi" | "December" | "Stoke" | | "Jenkinson" | "February" | "Arsenal" | ... | "Ivanovic" | "October" | "Chelsea" | | "Torres" | "October" | "Chelsea" | +--------------------------------------------+ So we can see that Stoke got 2 players sent off in December and Chelsea got 2 sent off in October. We can write the following query to return a result set which uses team and month as the grouping key i.e. we count how many paths there are which have the same team and month: START player = node:players('name:*') MATCH player-[:sent_off_in]-game-[:in_month]-month, game-[:in_match]-stats-[:stats]-player, stats-[:played_for]-team RETURN month.name, team.name, COUNT(player.name) AS numberOfReds ORDER BY numberOfReds DESC When we run that query we see the following results: +--------------------------------------------+ | month.name | team.name | numberOfReds | +--------------------------------------------+ | "December" | "Stoke" | 2 | | "October" | "Chelsea" | 2 | ... | "August" | "Stoke" | 1 | | "November" | "Tottenham" | 1 | | "December" | "Everton" | 1 | +--------------------------------------------+ This is all explained in more detail in the documentation but I thought it’d be interesting to write about it from the perspective of someone more used to writing SQL and trying to work out how to achieve the same thing in cypher.
February 19, 2013
by Mark Needham
· 27,391 Views
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XML->JSON->HashMap
Yes, it is long time since i posted… Was just trying to see how a XML can be converted to JSON and to HashMap. The situation is very imaginary. import java.io.File; import java.io.FileInputStream; import java.io.IOException; import java.io.InputStream; import java.util.ArrayList; import java.util.List; import java.util.Map; import net.sf.json.JSON; import net.sf.json.xml.XMLSerializer; import org.apache.commons.io.IOUtils; import org.codehaus.jackson.JsonGenerationException; import org.codehaus.jackson.map.JsonMappingException; import org.codehaus.jackson.map.ObjectMapper; import org.codehaus.jackson.type.TypeReference; public class XML2JSONConvertor { public static void main(String[] args) throws Exception { InputStream is = new FileInputStream(new File( “e:\\jagannathan\\personal\\java-projects\\secondtest.xml”)); String xml = IOUtils.toString(is); XMLSerializer xmlSerializer = new XMLSerializer(); JSON json = xmlSerializer.read(xml); System.out.println(json.toString(2)); printJSON(json.toString(2)); } public static void printJSON(String jsonString) { ObjectMapper mapper = new ObjectMapper(); try { Map jsonInMap = mapper.readValue(jsonString, new TypeReference>() { }); List keys = new ArrayList(jsonInMap.keySet()); for (String key : keys) { System.out.println(key + “: ” + jsonInMap.get(key)); } } catch (JsonGenerationException e) { e.printStackTrace(); } catch (JsonMappingException e) { e.printStackTrace(); } catch (IOException e) { e.printStackTrace(); } } } Dependencies net.sf.json-lib json-lib 2.4 jdk15 commons-io commons-io 2.3 compile xom xom 1.2.5 org.codehaus.jackson jackson-mapper-asl 1.9.0 The Input XML Jags Inc Jagan Male 24-jul Satya Male 24-apr The output 7 Feb, 2013 7:20:50 PM net.sf.json.xml.XMLSerializer getType INFO: Using default type string { “name”: “Jags Inc”, “employees”: [ { "name": "Jagan", "sex": "Male", "dob": "24-jul" }, { "name": "Satya", "sex": "Male", "dob": "24-apr" } ] } name: Jags Inc employees: [{name=Jagan, sex=Male, dob=24-jul}, {name=Satya, sex=Male, dob=24-apr}]
February 18, 2013
by Jagannathan Asokan
· 33,749 Views
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Building an Online-Recommendation Engine with MongoDB
once upon a time there was a munich pizza baker who developed a technique to beam pizza out of bright sunshine. he can produce more than a thousand pizzas per second and needs a channel to sell this amount of pizza and decides to build an online shop. mario’s initial idea is to sell pizzas, but now he is thinking about introduction of new product lines like beverages, salads and pasta. before we take a look to the validation of mario´s idea, lets take a short look at the existing online shop. mario’s online shop is based on mongodb , apache wicket and spring . mongodb is a document-oriented nosql-database . mongodb stores records not in tables as a relational database but in bson documents, which is a binary version of json (java script object notation) and very similar to the object structure in mario’s application. the usage of mongodb makes his development easier and deployment faster. the figure shows a json document which is very similar to a java object: a json document property with the according value corresponds to the java object property with the appropriate value. you can add or remove properties in your java object and this will automatically change your database schema. so there is no need to put your java object model into a relational schema via hibernate. mario also decided to build his online shop only with open-source technologies like apache wicket and spring. wicket is a very common lightweight component-based web application framework and it is closely patterned after stateful gui frameworks such as javafx . the spring framework is an open source application framework and inversion of control container for the java platform and does not impose any specific programming model. spring has become popular in the java community as an alternative to, replacement for, or even addition to the enterprise javabean (ejb) model. because of this architecture mario is able to deploy its application in a lightweight application server like tomcat or jetty . this figure shows the system landscape of mario. mario has two major system on the lefthand site there is his online shop and on the righthand site there is ‘pas’ a famous billing system. in the middle is hadoop that connects both systems together. in the business world an application normally does not stand alone. in most cases an application must communicate with others. the lean architecture of marios online shop enables him to connect the billing system ‘pas’ to his online shop. spring for apache hadoop provides this integration between the two systems online shop and ‘pas’. hadoop supports data-intensive distributed applications and implements a computational paradigm named mapreduce, where the computation is divided into many small fragments, each of them may be executed or re-executed on any node in the cluster of commodity hardware. mario uses hadoop as an etl layer that enables him to transfer gigabytes of order information into the billing system. in this case hadoop makes it possible for a financial controller to verify if all orders were billed correctly. in addition to the online shop feature mario has a real-time sales dashboard that enables him to track his sales in real time. the dashboard displays daily and monthly sales statistics for each pizza and contains a map with the geographical overview of customer activity and competitor locations. here is a walkthrough of the shop : now lets talk about mario’s incredible new idea : mario wants to sell even more pizza! and other products as well. mario decides to use lean startup methods in order to test the possible introduction of new product lines and plans an experiment to validate his new idea using a scientific approach and pure facts instead of hunches. mario´s core assumption is that customers wants to buy other products than pizza – drinks, salads and pasta. furthermore he is worried about pricing. mario contacts all customers to complete a survey and provides an incentive for the participation, a free pizza to every customer who responds to the survey. the result of the survey validated mario’s assumption – customers want to buy beverages, salads and pasta. but he also found out that his customers are willing to pay higher prices for high-quality products and that they simply love his easy shopping flow. currently a pizza order can be completed with three clicks only, so there is new riskiest assumption to validate: will a more complex shopping flow affect his sales? the figures shows a validation board. a validation board is a deceptively simple tool for testing out product ideas. furthermore a validation board tracks pivots which follows from customer feedback. mario decides to introduce beverages, salads and pasta product lines and thinks about a possibility, how he can handle the extension of the product line without destroying the easy shopping flow. that’s why mario thinks a recommendation engine is the right way for him. panels for recommendations can be integrated in the online shop without changing the shopping flow. mario hired a statistician to help him implement a recommender system for his online shop for better cross-selling. he also defined new measurement points to validate his new idea . therefore he tracks the conversion rate of orders as well as cross-selling rates and every event in the online shop is already tracked in realtime. so mario can very easily perform further experiments in order to verify more assumptions. follow the blog to see how the story continues or come to mongodb usergroup meetup in munich , february 20, 2013 or mongodb days in berlin , february 26, 2013 to get a live presentation. our talk sheds light on how to build an online recommendation engine based on mongodb and apache mahout. we’ll show which recommenders must be built to reach mario’s goal and how these can be integrated in mario’s shop infrastructure.
February 17, 2013
by Comsysto Gmbh
· 8,609 Views
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Better explaining the CAP Theorem
today, i thought a lot about how to examine different databases. choosing a database is often a daunting task. there's a lot of confusion, a 'theorem', and more than all, the immortal proverb 'not one size fits all'. as if it helps. one of the first things that you realize, when examining nosql distributed databases (and how could you not)is that these days databases are like cars: they're all good. old fashioned sql databases can scale in and out, horizontally sharded over several machines to achieve high availability. nosql systems claim to be consistent. what difference then does it make what database would you choose? the availability and consistency that i mentioned comes, of course, from the misunderstood cap theorem , that - so people say - states that you can only choose 2 out of the 3 consistency: every read would get you the most recent write availability: every node (if not failed) always executes queries partition-tolerance: even if the connections between nodes are down, the other two (a & c) promises, are kept. usually its depicted in a nicely equilaterl triangle, as this one from ofirm : there's a nice proof and explanation of it in this 4 minute video here . but if we think about it, and also see some of brewer's (the theorem author) later remarks , we'll see that the 2 out of 3 is really 1 out of 2: it's really just a vs c! and this is simply because: availability is achieved by replicating the data across different machines consistency is achieved by updating several nodes before allowing further reads total partitioning, meaning failure of part of the system is rare. however, we could look at a delay, a latency, of the update between nodes, as a temporary partitioning . it will then cause a temporary decision between a and c: on systems that allow reads before updating all the nodes, we will get high availability on systems that lock all the nodes before allowing reads, we will get consistency that's it! and since this decision is temporary, it exists only for the duration of the delay, some may say that we are really contrasting latency (another word for availability) against consistency. by the way, there's no distributed system that wants to live with "paritioning" - if it does, it's not distributed. that is why putting sql in this triangle may lead to confusion.
February 17, 2013
by Lior Messinger
· 139,699 Views · 18 Likes
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CPU Cache Flushing Fallacy
Even from highly experienced technologists I often hear talk about how certain operations cause a CPU cache to "flush". This seems to be illustrating a very common fallacy about how CPU caches work, and how the cache sub-system interacts with the execution cores. In this article I will attempt to explain the function CPU caches fulfil, and how the cores, which execute our programs of instructions, interact with them. For a concrete example I will dive into one of the latest Intel x86 server CPUs. Other CPUs use similar techniques to achieve the same ends. Most modern systems that execute our programs are shared-memory multi-processor systems in design. A shared-memory system has a single memory resource that is accessed by 2 or more independent CPU cores. Latency to main memory is highly variable from 10s to 100s of nanoseconds. Within 100ns it is possible for a 3.0GHz CPU to process up to 1200 instructions. Each Sandy Bridge core is capable of retiring up to 4 instructions-per-cycle (IPC) in parallel. CPUs employ cache sub-systems to hide this latency and allow them to exercise their huge capacity to process instructions. Some of these caches are small, very fast, and local to each core; others are slower, larger, and shared across cores. Together with registers and main-memory, these caches make up our non-persistent memory hierarchy. Next time you are developing an important algorithm, try pondering that a cache-miss is a lost opportunity to have executed ~500 CPU instructions! This is for a single-socket system, on a multi-socket system you can effectively double the lost opportunity as memory requests cross socket interconnects. Memory Hierarchy Figure 1. For the circa 2012 Sandy Bridge E class servers our memory hierarchy can be decomposed as follows: Registers: Within each core are separate register files containing 160 entries for integers and 144 floating point numbers. These registers are accessible within a single cycle and constitute the fastest memory available to our execution cores. Compilers will allocate our local variables and function arguments to these registers. When hyperthreading is enabled these registers are shared between the co-located hyperthreads. Memory Ordering Buffers (MOB): The MOB is comprised of a 64-entry load and 36-entry store buffer. These buffers are used to track in-flight operations while waiting on the cache sub-system. The store buffer is a fully associative queue that can be searched for existing store operations, which have been queued when waiting on the L1 cache. These buffers enable our fast processors to run asynchronously while data is transferred to and from the cache sub-system. When the processor issues asynchronous reads and writes then the results can come back out-of-order. The MOB is used to disambiguate the ordering for compliance to the published memory model. Level 1 Cache: The L1 is a core-local cache split into separate 32K data and 32K instruction caches. Access time is 3 cycles and can be hidden as instructions are pipelined by the core for data already in the L1 cache. Level 2 Cache: The L2 cache is a core-local cache designed to buffer access between the L1 and the shared L3 cache. The L2 cache is 256K in size and acts as an effective queue of memory accesses between the L1 and L3. L2 contains both data and instructions. L2 access latency is 12 cycles. Level 3 Cache: The L3 cache is shared across all cores within a socket. The L3 is split into 2MB segments each connected to a ring-bus network on the socket. Each core is also connected to this ring-bus. Addresses are hashed to segments for greater throughput. Latency can be up to 38 cycles depending on cache size. Cache size can be up to 20MB depending on the number of segments, with each additional hop around the ring taking an additional cycle. The L3 cache is inclusive of all data in the L1 and L2 for each core on the same socket. This inclusiveness, at the cost of space, allows the L3 cache to intercept requests thus removing the burden from private core-local L1 & L2 caches. Main Memory: DRAM channels are connected to each socket with an average latency of ~65ns for socket local access on a full cache-miss. This is however extremely variable, being much less for subsequent accesses to columns in the same row buffer, through to significantly more when queuing effects and memory refresh cycles conflict. 4 memory channels are aggregated together on each socket for throughput, and to hide latency via pipelining on the independent memory channels. NUMA: In a multi-socket server we have non-uniform memory access. It is non-uniform because the required memory maybe on a remote socket having an additional 40ns hop across the QPI bus. Sandy Bridge is a major step forward for 2-socket systems over Westmere and Nehalem. With Sandy Bridge the QPI limit has been raised from 6.4GT/s to 8.0GT/s, and two lanes can be aggregated thus eliminating the bottleneck of the previous systems. For Nehalem and Westmere the QPI link is only capable of ~40% the bandwidth that could be delivered by the memory controller for an individual socket. This limitation made accessing remote memory a choke point. In addition, the QPI link can now forward pre-fetch requests which previous generations could not. Associativity Levels Caches are effectively hardware based hash tables. The hash function is usually a simple masking of some low-order bits for cache indexing. Hash tables need some means to handle a collision for the same slot. The associativity level is the number of slots, also known as ways or sets, which can be used to hold a hashed version of an address. Having more levels of associativity is a trade off between storing more data vs. power requirements and time to search each of the ways. For Sandy Bridge the L1 and L2 are 8-way and the L3 is 12-way associative. Cache Coherence With some caches being local to cores, we need a means of keeping them coherent so all cores can have a consistent view of memory. The cache sub-system is considered the "source of truth" for mainstream systems. If memory is fetched from the cache it is never stale; the cache is the master copy when data exists in both the cache and main-memory. This style of memory management is known as write-back whereby data in the cache is only written back to main-memory when the cache-line is evicted because a new line is taking its place. An x86 cache works on blocks of data that are 64-bytes in size, known as a cache-line. Other processors can use a different size for the cache-line. A larger cache-line size reduces effective latency at the expense of increased bandwidth requirements. To keep the caches coherent the cache controller tracks the state of each cache-line as being in one of a finite number of states. The protocol Intel employs for this is MESIF, AMD employs a variant know as MOESI. Under the MESIF protocol each cache-line can be in 1 of the 5 following states: Modified: Indicates the cache-line is dirty and must be written back to memory at a later stage. When written back to main-memory the state transitions to Exclusive. Exclusive: Indicates the cache-line is held exclusively and that it matches main-memory. When written to, the state then transitions to Modified. To achieve this state a Request-For-Ownership (RFO) message is sent which involves a read plus an invalidate broadcast to all other copies. Shared: Indicates a clean copy of a cache-line that matches main-memory. Invalid: Indicates an unused cache-line. Forward: Indicates a specialised version of the shared state i.e. this is the designated cache which should respond to other caches in a NUMA system. To transition from one state to another, a series of messages are sent between the caches to effect state changes. Previous to Nehalem for Intel, and Opteron for AMD, this cache coherence traffic between sockets had to share the memory bus which greatly limited scalability. These days the memory controller traffic is on a separate bus. The Intel QPI, and AMD HyperTransport, buses are used for cache coherence between sockets. The cache controller exists as a module within each L3 cache segment that is connected to the on-socket ring-bus network. Each core, L3 cache segment, QPI controller, memory controller, and integrated graphics sub-system are connected to this ring-bus. The ring is made up of 4 independent lanes for: request, snoop, acknowledge, and 32-bytes data per cycle. The L3 cache is inclusive in that any cache-line held in the L1 or L2 caches is also held in the L3. This provides for rapid identification of the core containing a modified line when snooping for changes. The cache controller for the L3 segment keeps track of which core could have a modified version of a cache-line it owns. If a core wants to read some memory, and it does not have it in a Shared, Exclusive, or Modified state; then it must make a read on the ring bus. It will then either be read from main-memory if not in the cache sub-systems, or read from L3 if clean, or snooped from another core if Modified. In any case the read will never return a stale copy from the cache sub-system, it is guaranteed to be coherent. Concurrent Programming If our caches are always coherent then why do we worry about visibility when writing concurrent programs? This is because within our cores, in their quest for ever greater performance, data modifications can appear out-of-order to other threads. There are 2 major reasons for this. Firstly, our compilers can generate programs that store variables in registers for relatively long periods of time for performance reasons, e.g. variables used repeatedly within a loop. If we need these variables to be visible across cores then the updates must not be register allocated. This is achieved in C by qualifying a variable as "volatile". Beware that C/C++ volatile is inadequate for telling the compiler to order other instructions. For this you need fences/barriers. The second major issue with ordering we have to be aware of is a thread could write a variable and then, if it reads it shortly after, could see the value in its store buffer which may be older than the latest value in the cache sub-system. This is never an issue for algorithms following the Single Writer Principle but is an issue for the likes of the Dekker and Peterson lock algorithms. To overcome this issue, and ensure the latest value is observed, the thread must wait for the store buffer to drain on that core. This can be achieved by issuing a fence instruction. The write of a volatile variable in Java, in addition to never being register allocated, is accompanied by a full fence instruction. This fence instruction on x86 has a significant performance impact by preventing progress on the issuing thread until the store buffer is drained. Fences on other processors can have more efficient implementations that simply put a marker in the store buffer for the search boundary, e.g. the Azul Vega does this. If you want to ensure memory ordering across Java threads when following the Single Writer Principle, and avoid the store fence, it is possible by using the j.u.c.Atomic(Int|Long|Reference).lazySet() method, as opposed to setting a volatile variable. The Fallacy Returning to the fallacy of "flushing the cache" as part of a concurrent algorithm. I think we can safely say that we never "flush" the CPU cache within our user space programs. I believe the source of this fallacy is the need to flush, mark or drain to a point, the store buffer for some classes of concurrent algorithms so the latest value can be observed on a subsequent load operation. For this we require a memory ordering fence and not a cache flush. Another possible source of this fallacy is that L1 caches, or the TLB, may need to be flushed based on address indexing policy on a context switch. ARM, previous to ARMv6, did not use address space tags on TLB entries thus requiring the whole L1 cache to be flushed on a context switch. Many processors require the L1 instruction cache to be flushed for similar reasons, in many cases this is simply because instruction caches are not required to be kept coherent. The bottom line is, context switching is expensive and a bit off topic, so in addition to the cache pollution of the L2, a context switch can also cause the TLB and/or L1 caches to require a flush. Intel x86 processors require only a TLB flush on context switch.
February 15, 2013
by Martin Thompson
· 11,734 Views · 3 Likes
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The Reality of a Developer's Life — In GIFs, of Course
Want to check out what life is like as a developer? Check out this post for a series of GIFs demonstrating the dev life.
February 15, 2013
by Alex Soto
· 501,127 Views · 17 Likes
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Using awk and Friends with Hadoop
imagine you have a csv file that you want to manipulate. here’s a sample file we can play with: lopez,charlie,2002,11,21 parker,ward,1995,04,08 henderson,russell,2007,10,01 our goal is to transform this into the following form by combining the last three columns: lopez,charlie,20021121 parker,ward,19950408 henderson,russell,20071001 in linux this would take all of two seconds (excuse the awkward awk command): shell$ awk -f"," '{ print $1","$2","$3$4$5 }' people.txt what if you wanted to quickly do the same in hdfs - and let’s assume you want to write the results back to hdfs. one approach would be to use the hdfs cli to stream the inputs into awk, and stream the awk output back into hdfs. you could do this with the hdfs cat and put - options (note that adding a hyphen after put instructs the put command to stream data from standard input to hdfs): shell$ hadoop fs -cat people.txt | awk -f"," '{ print $1","$2","$3$4$5 }' | hadoop fs -put - people-coalesed.txt btw, if your input and output files are lzop-compressed then this command would work: shell$ hadoop fs -cat people.txt.lzo | lzop -dc | awk -f"," '{ print $1","$2","$3$4$5 }' | \ lzop -c | hadoop fs -put - people-coalesed.txt.lzo this is great if your file isn’t too large, but if it’s multiple gigabytes in length then you probably want to harness the power of mapreduce to get this done in a jiffy! the words “in a jiffy” and “mapreduce” aren’t commonly used together, so what do we do? well you could crack open pig or hive and write some custom user-defined functions, but this means you end up in java which we want to avoid. hadoop streaming comes to the rescue in these situations. let’s first create our awk script which will be executed: shell$ cat people.awk #!/bin/awk -f begin { fs = "," } { print $1","$2","$3$4$5 } in linux, if you make this awk script executable, you could execute is as follows: shell$ ./people.awk people.txt in mapreduce-land we don’t need to join data in this particular example, so we don’t need to run any reducers. call your awk script from mappers via hadoop streaming with this command: shell$ hadoop_home=/usr/lib/hadoop shell$ ${hadoop_home}/bin/hadoop \ jar ${hadoop_home}/contrib/streaming/*.jar \ -d mapreduce.job.reduces=0 \ -d mapred.reduce.tasks=0 \ -input people.txt \ -output people-coalesed \ -mapper people.awk \ -file people.awk a few options in the hadoop streaming command are worth examining: finally - to get lzo into the picture you need to add -inputformat , -d mapred.output.compress and -d mapred.output.compression.codec arguments: shell$ hadoop_home=/usr/lib/hadoop shell$ ${hadoop_home}/bin/hadoop \ jar ${hadoop_home}/contrib/streaming/*.jar \ -d mapreduce.job.reduces=0 \ -d mapred.reduce.tasks=0 \ -d mapred.output.compress=true \ -d stream.map.input.ignorekey=true \ -d mapred.output.compression.codec=com.hadoop.compression.lzo.lzopcodec \ -inputformat com.hadoop.mapred.deprecatedlzotextinputformat \ -input people.txt.lzo \ -output people-coalesed \ -mapper people.awk \ -file people.awk
February 14, 2013
by Alex Holmes
· 13,238 Views · 1 Like
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Introduction to JCache JSR 107
Resin has supported caching, session replication (another form of caching), and http proxy caching in cluster environments for over ten years. When you use Resin caching, you are using the same platform that has the speed and scalability of custom services written in C like NginX with the usability of Java, and the industry platform Java EE. JCache JSR 107 is a distributed cache that has a similar interface to the HashMap that you know and love. To be more specific, the Cache object in JCache looks like a java.util.ConncurrentHashMap. In addition, JCache JSR 107 defines integration with CDI (as well as Spring and Guice). You can decorate services with interceptors that apply caching to the services just by defining annotations. Resin 4 has support for JCache, and JCache support is required for Java EE 7. Let's look at a small example to see how easy is to get started with JCache. package hello.world; import javax.cache.Cache; import javax.cache.CacheBuilder; import javax.cache.CacheManager; import javax.cache.Caching; ... @WebServlet("/HelloServlet") public class HelloServlet extends HttpServlet { Cache cache; public Cache cache() { if (cache == null) { //building a cache CacheManager manager = Caching.getCacheManager("cacheManagerHello"); CacheBuilder builder = manager.createCacheBuilder("a"); cache = builder.build(); } return cache; } protected void doGet(HttpServletRequest request, HttpServletResponse response) throws ServletException, IOException { response.setContentType("text/html"); response.getWriter().append(" "); String helloMessage = cache().get("hello message"); if (helloMessage == null) { helloMessage = new StringBuilder(20) .append("Hello World ! ") .append(System.currentTimeMillis()).toString(); cache().put("hello message", helloMessage); // <-------------- putting results in the cache } response.getWriter().append(helloMessage); response.getWriter().append(" "); } } The above works out fairly well, but what if we want to periodically change the helloMessage. Let's say we get 2,000 requests a second, but every 10 seconds or so we would like to regenerate the helloMessage. The message might be: Hello World ! 1358979745996 Later we would want it to change. If we wanted it to change every 10 seconds after it was last accessed, we would do this: cache = builder.setExpiry(ExpiryType.ACCESSED, new Duration(TimeUnit.SECONDS, 10)).build(); For this example, we want to change it every 10 seconds after is was last modified. We would set up the timeout on the creation as follows: cache = builder.setExpiry(ExpiryType.MODIFIED, new Duration(TimeUnit.SECONDS, 10)).build(); This would go right in the cache method we defined earlier. public Cache cache() { if (cache == null) { CacheManager manager = Caching.getCacheManager("cacheManagerHello"); CacheBuilder builder = manager.createCacheBuilder("b"); cache = builder.setExpiry(ExpiryType.MODIFIED, new Duration(TimeUnit.SECONDS, 10)).build(); } return cache; } Resin's JCache implementation is built on top Resin distributed cache architecture. You get replication, and data redundancy built in. Bill Digman is a Java EE / Servlet enthusiast and Open Source enthusiast who loves working with Caucho's Resin Servlet Container, a Java EE Web Profile Servlet Container. Caucho's Resin OpenSource Servlet Container Java EE Web Profile Servlet Container Caucho's Resin 4.0 JCache blog post
February 13, 2013
by Bill Digman
· 49,176 Views · 1 Like
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Synchronising Multithreaded Integration Tests
Testing threads is hard, very hard and this makes writing good integration tests for multithreaded systems under test... hard. This is because in JUnit there's no built in synchronisation between the test code, the object under test and any threads. This means that problems usually arise when you have to write a test for a method that creates and runs a thread. One of the most common scenarios in this domain is in making a call to a method under test, which starts a new thread running before returning. At some point in the future when the thread's job is done you need assert that everything went well. Examples of this scenario could include asynchronously reading data from a socket or carrying out a long and complex set of operations on a database. For example, the ThreadWrapper class below contains a single public method: doWork(). Calling doWork() sets the ball rolling and at some point in the future, at the discretion of the JVM, a thread runs adding data to a database. public class ThreadWrapper { /** * Start the thread running so that it does some work. */ public void doWork() { Thread thread = new Thread() { /** * Run method adding data to a fictitious database */ @Override public void run() { System.out.println("Start of the thread"); addDataToDB(); System.out.println("End of the thread method"); } private void addDataToDB() { // Dummy Code... try { Thread.sleep(4000); } catch (InterruptedException e) { e.printStackTrace(); } } }; thread.start(); System.out.println("Off and running..."); } } A straightforward test for this code would be to call the doWork() method and then check the database for the result. The problem is that, owing to the use of a thread, there's no co-ordination between the object under test, the test and the thread. A common way of achieving some co-ordination when writing this kind of test is to put some kind of delay in between the call to the method under test and checking the results in the database as demonstrated below: public class ThreadWrapperTest { @Test public void testDoWork() throws InterruptedException { ThreadWrapper instance = new ThreadWrapper(); instance.doWork(); Thread.sleep(10000); boolean result = getResultFromDatabase(); assertTrue(result); } /** * Dummy database method - just return true */ private boolean getResultFromDatabase() { return true; } } In the code above there is a simple Thread.sleep(10000) between two method calls. This technique has the benefit of being incredabile simple; however it's also very risky. This is because it introduces a race condition between the test and the worker thread as the JVM makes no guarantees about when threads will run. Often it'll work on a developer's machine only to fail consistently on the build machine. Even if it does work on the build machine it atificially lengthens the duration of the test; remember that quick builds are important. The only sure way of getting this right is to synchronise the two different threads and one technique for doing this is to inject a simple CountDownLatch into the instance under test. In the example below I've modified the ThreadWrapper class's doWork() method adding the CountDownLatch as an argument. public class ThreadWrapper { /** * Start the thread running so that it does some work. */ public void doWork(final CountDownLatch latch) { Thread thread = new Thread() { /** * Run method adding data to a fictitious database */ @Override public void run() { System.out.println("Start of the thread"); addDataToDB(); System.out.println("End of the thread method"); countDown(); } private void addDataToDB() { try { Thread.sleep(4000); } catch (InterruptedException e) { e.printStackTrace(); } } private void countDown() { if (isNotNull(latch)) { latch.countDown(); } } private boolean isNotNull(Object obj) { return latch != null; } }; thread.start(); System.out.println("Off and running..."); } } he Javadoc API describes a count down latch as: A synchronization aid that allows one or more threads to wait until a set of operations being performed in other threads completes. A CountDownLatch is initialized with a given count. The await methods block until the current count reaches zero due to invocations of the countDown() method, after which all waiting threads are released and any subsequent invocations of await return immediately. This is a one-shot phenomenon -- the count cannot be reset. If you need a version that resets the count, consider using a CyclicBarrier. A CountDownLatch is a versatile synchronization tool and can be used for a number of purposes. A CountDownLatch initialized with a count of one serves as a simple on/off latch, or gate: all threads invoking await wait at the gate until it is opened by a thread invoking countDown(). A CountDownLatchinitialized to N can be used to make one thread wait until N threads have completed some action, or some action has been completed N times. A useful property of a CountDownLatch is that it doesn't require that threads calling countDown wait for the count to reach zero before proceeding, it simply prevents any thread from proceeding past an await until all threads could pass. The idea here is that the test code will never check the database for the results until the run() method of the worker thread has called latch.countdown(). This is because the test code thread is blocking at the call to latch.await(). latch.countdown() decrements latch's count and once this is zero the blocking call the latch.await() returns and the test code continues executing, safe in the knowledge that any results which should be in the database, are in the database. The test can then retrieve these results and make a valid assertion. Obviously, the code above merely fakes the database connection and operations. The thing is you may not want to, or need to, inject a CountDownLatch directly into your code; after all it's not used in production and it doesn't look particularly clean or elegant. One quick way around this is to simply make the doWork(CountDownLatch latch) method package private and expose it through a public doWork() method. public class ThreadWrapper { /** * Start the thread running so that it does some work. */ public void doWork() { doWork(null); } @VisibleForTesting void doWork(final CountDownLatch latch) { Thread thread = new Thread() { /** * Run method adding data to a fictitious database */ @Override public void run() { System.out.println("Start of the thread"); addDataToDB(); System.out.println("End of the thread method"); countDown(); } private void addDataToDB() { try { Thread.sleep(4000); } catch (InterruptedException e) { e.printStackTrace(); } } private void countDown() { if (isNotNull(latch)) { latch.countDown(); } } private boolean isNotNull(Object obj) { return latch != null; } }; thread.start(); System.out.println("Off and running..."); } } The code above uses Google's Guava @VisibleForTesting annotation to tell us that the doWork(CountDownLatch latch) method visibility has been relaxed slightly for testing purposes. Now I realise that making a method call package private for testing purposes in highly controversial; some people hate the idea, whilst others include it everywhere. I could write a whole blog on this subject (and may do one day), but for me it should be used judiciously, when there's no other choice, for example when you're writing characterisation tests for legacy code. If possible it should be avoided, but never ruled out. After all tested code is better than untested code. With this in mind the next iteration of ThreadWrapper designs out the need for a method marked as @VisibleForTesting together with the need to inject a CountDownLatch into your production code. The idea here is to use the Strategy Pattern and separate the Runnable implementation from the Thread. Hence, we have a very simple ThreadWrapper public class ThreadWrapper { /** * Start the thread running so that it does some work. */ public void doWork(Runnable job) { Thread thread = new Thread(job); thread.start(); System.out.println("Off and running..."); } } and a separate job: public class DatabaseJob implements Runnable { /** * Run method adding data to a fictitious database */ @Override public void run() { System.out.println("Start of the thread"); addDataToDB(); System.out.println("End of the thread method"); } private void addDataToDB() { try { Thread.sleep(4000); } catch (InterruptedException e) { e.printStackTrace(); } } } You'll notice that the DatabaseJob class doesn't use a CountDownLatch. How is it synchronised? The answer lies in the test code below... public class ThreadWrapperTest { @Test public void testDoWork() throws InterruptedException { ThreadWrapper instance = new ThreadWrapper(); CountDownLatch latch = new CountDownLatch(1); DatabaseJobTester tester = new DatabaseJobTester(latch); instance.doWork(tester); latch.await(); boolean result = getResultFromDatabase(); assertTrue(result); } /** * Dummy database method - just return true */ private boolean getResultFromDatabase() { return true; } private class DatabaseJobTester extends DatabaseJob { private final CountDownLatch latch; public DatabaseJobTester(CountDownLatch latch) { super(); this.latch = latch; } @Override public void run() { super.run(); latch.countDown(); } } } The test code above contains an inner class DatabaseJobTester, which extends DatabaseJob. In this class the run() method has been overridden to include a call to latch.countDown() after our fake database has been updated via the call to super.run(). This works because the test passes a DatabaseJobTester instance to the doWork(Runnable job) method adding in the required thread testing capability. The idea of sub-classing objects under test is something I've mentioned before in one of my blogs on testing techniques and is a really powerful technique. So, to conclude: Testing threads is hard. Testing anonymous inner classes is almost impossible. Using Thead.sleep(...) is a risky idea and should be avoided. You can refactor out these problems using the Strategy Pattern. Programming is the Art of Making the Right Decision ...and that relaxing a method's visibility for testing may or may not be a good idea, but more on that later... The code above is available on Github in the captain debug repository (git://github.com/roghughe/captaindebug.git) under the unit-testing-threads project.
February 13, 2013
by Roger Hughes
· 14,152 Views · 12 Likes
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How to Return Dictionary As a Result From a LINQ Query in C#?
This article will provide a code snippet and explains how to return Dictionary as result from a LINQ Query in C#. There are times when you want to retrieve only the ID(distinct) and the name from the database table using LINQ . In scenarios like this , one can use the ToDictionary method to place the necessary properties to the dictionary and return them. Below is a sample sourecode demonstrating the usage of ToDictionary method in LINQ Query public class BlockbusterMovie { public string Name { get; set; } public int ID { get; set; } } public class BlockbusterMovies : List { public BlockbusterMovies() { Add(new BlockbusterMovie { Name = "Vishwaroopam", ID = 1 }); Add(new BlockbusterMovie { Name = "Endhiran", ID = 2 }); Add(new BlockbusterMovie { Name = "Thuppaki", ID = 3 }); Add(new BlockbusterMovie { Name = "Mankatha", ID = 4 }); } } The BlockbusterMovies class has the collection of movies which is used in the below code snippet to return the dictionary based on the ID and Name. private void Form1_Load(object sender, EventArgs e) { List movies = new BlockbusterMovies(); var LstMovies = movies.ToDictionary(Field => Field.ID, mc => mc.Name); }
February 13, 2013
by Senthil Kumar
· 53,520 Views
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How to create/restore a slave using GTID replication in MySQL 5.6
This post comes from Miguel Angel Nieto at the MySQL Performance Blog. MySQL 5.6 is GA! Now we have new things to play with and in my personal opinion the most interesting one is the new Global Transaction ID (GTID) support in replication. This post is not an explanation of what is GTID and how it works internally because there are many documents about that: http://dev.mysql.com/doc/refman/5.6/en/replication-gtids-concepts.html One thing that worths to mention is that if you want GTID support log_slave_updates will need to be enabled in slave server and the performance impact should be taken in account. Anyway, this post tends to be more practical, we will see how to create/restore new slaves from a master using GTID. How to set up a new slave The first thing that we need to know is that now Binary Logs and Position are not needed anymore with GTID enabled. Instead we need to know in which GTID is the master and set it on the slave. MySQL keeps two global variables with GTID numbers on it: gtid_executed: it contains a representation of the set of all transaction logged in the binary log gtid_purged: it contains a representation of the set of all transactions deleted from the binary log So now, the process is the following: take a backup from the master and store the value of gtid_executed restore the backup on the slave and set gtid_purged with the value of gtid_executed from the master The new mysqldump can do those tasks for us. Let’s see an example of how to take a backup from the master and restore it on the slave to set up a new replication server. master > show global variables like 'gtid_executed'; +---------------+-------------------------------------------+ | Variable_name | Value | +---------------+-------------------------------------------+ | gtid_executed | 9a511b7b-7059-11e2-9a24-08002762b8af:1-13 | +---------------+-------------------------------------------+ master > show global variables like 'gtid_purged'; +---------------+------------------------------------------+ | Variable_name | Value | +---------------+------------------------------------------+ | gtid_purged | 9a511b7b-7059-11e2-9a24-08002762b8af:1-2 | +---------------+------------------------------------------+ Now we take a backup with mysqldump from the master: # mysqldump --all-databases --single-transaction --triggers --routines --host=127.0.0.1 --port=18675 --user=msandbox --password=msandbox > dump.sql It will contain the following line: # grep PURGED dump.sql SET @@GLOBAL.GTID_PURGED='9a511b7b-7059-11e2-9a24-08002762b8af:1-13'; Therefore during the dump recover process on the slave it will set GTID_PURGED to the GTID_EXECUTED value from the master. So now, we just need to recover the dump and start the replication: slave1 > show global variables like 'gtid_executed'; +---------------+-------+ | Variable_name | Value | +---------------+-------+ | gtid_executed | | +---------------+-------+ slave1 > show global variables like 'gtid_purged'; +---------------+-------+ | Variable_name | Value | +---------------+-------+ | gtid_purged | | +---------------+-------+ slave1 > slave1> source test.sql; [...] slave1 > show global variables like 'gtid_executed'; +---------------+-------------------------------------------+ | Variable_name | Value | +---------------+-------------------------------------------+ | gtid_executed | 9a511b7b-7059-11e2-9a24-08002762b8af:1-13 | +---------------+-------------------------------------------+ slave1 > show global variables like 'gtid_purged'; +---------------+-------------------------------------------+ | Variable_name | Value | +---------------+-------------------------------------------+ | gtid_purged | 9a511b7b-7059-11e2-9a24-08002762b8af:1-13 | +---------------+-------------------------------------------+ The last step is to configure the slave using the auto-configuration method of GTID: slave1 > CHANGE MASTER TO MASTER_HOST="127.0.0.1", MASTER_USER="msandbox", MASTER_PASSWORD="msandbox", MASTER_PORT=18675, MASTER_AUTO_POSITION = 1; How to restore a slave in a bad and fast way Let’s imagine that our slave has been down for several days and the binary logs from the master have been purged. This is the error we are going to get: Slave_IO_Running: No Slave_SQL_Running: Yes Last_IO_Error: Got fatal error 1236 from master when reading data from binary log: 'The slave is connecting using CHANGE MASTER TO MASTER_AUTO_POSITION = 1, but the master has purged binary logs containing GTIDs that the slave requires.' So, let’s try to solve it. First we have the bad and fast way, that is, point to another GTID that the master has in the binary logs. First, we get the GTID_EXECUTED from the master: master > show global variables like 'GTID_EXECUTED'; +---------------+-------------------------------------------+ | Variable_name | Value | +---------------+-------------------------------------------+ | gtid_executed | 9a511b7b-7059-11e2-9a24-08002762b8af:1-14 | +---------------+-------------------------------------------+ And we set it on the slave: slave> set global GTID_EXECUTED="9a511b7b-7059-11e2-9a24-08002762b8af:1-14" ERROR 1238 (HY000): Variable 'gtid_executed' is a read only variable Error! Remember, we get the GTID_EXECUTED from the master and set is as GTID_PURGED on the slave. slave1 > set global GTID_PURGED="9a511b7b-7059-11e2-9a24-08002762b8af:1-14"; ERROR 1840 (HY000): GTID_PURGED can only be set when GTID_EXECUTED is empty. Error again, GTID_EXECUTED should be empty before changing GTID_PURGED manually but we can’t change it with SET because is a read only variable. The only way to change it is with reset master (yes, on a slave server): slave1> reset master; slave1 > show global variables like 'GTID_EXECUTED'; +---------------+-------+ | Variable_name | Value | +---------------+-------+ | gtid_executed | | +---------------+-------+ slave1 > set global GTID_PURGED="9a511b7b-7059-11e2-9a24-08002762b8af:1-14"; slave1> start slave io_thread; slave1> show slave status\G [...] Slave_IO_Running: Yes Slave_SQL_Running: Yes [...] Now, if you don’t get any error like primary/unique key duplication then you can run the pt-table-checksum and pt-table-sync. How to restore a slave in a good and slow way The good way is mysqldump again. We take a dump from the master like we saw before and try to restore it on the slave: slave1 [localhost] {msandbox} ((none)) > source test.sql; [...] ERROR 1840 (HY000): GTID_PURGED can only be set when GTID_EXECUTED is empty. [...] Wop! It is important to mention that these kind of error messages can dissapear on the shell buffer because the restore of the dump will continue. Be cautious. Same problem again so same solution too: slave1> reset master; slave1> source test.sql; slave1> start slave; slave1> show slave status\G [...] Slave_IO_Running: Yes Slave_SQL_Running: Yes [...] Conclusion With the new GTID we need to change our minds. Now binary log and position is not something we need to take in account, gtid_executed and gtid_purged are our new friends. Xtrabackup still doesn’t support it but we are working on it. I will update this post and create a one when we publish a xtrabackup version with full support of GTID.
February 11, 2013
by Peter Zaitsev
· 18,876 Views
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Building SOLID Databases: Liskov Substitution Weirdness
if the open/closed principle starts looking a little strange as applied to object-relational design, the liskov substitution principle applies in almost the exactly opposite way in the database as in the applications. the reason become only clear once exploring this principle, seeing the limits on it, and investigating the realities of database development. the canonical examples of lsp violations in application programs do not violate the principle at all in the database. as always the reason is that databases model different things than applications and so their constraints are different. the formal definition of the liskov substitution principle let be a property provable about objects of type . then should be provable for objects of type where is a subtype of . the liskov substitution principle in application programming the liskov substitution principle is important because it ensures provability of state and behavior for subtypes. this means you have to look at the full constraints that operate on a class method (excluding constructors i think, which operate under a different set of constraints). in general for the lsp to be satisfied, preconstraints cannot be weakened, post-constraints cannot be strengthened, and invariants must be maintained not only in inherited methods but also non-inherited ones (this is broad enough to include the history constraint). the typical problem this addresses is the square-rectangle problem, which is a well-understood problem in object-oriented programming, and one which calls into question certain promises of oop as a whole. the square-rectangle problem establishes a problem of irreducible complexity in object-oriented programming. the typical examples are either "a square is-a rectangle" or "a circle is-a elipse." mathematically both of these statements are true but implementing them in an object-oriented framework adds a great deal of complexity. the reason is that the mathematical hierarchy has no concept of a change of state which preserves invariances, and so the object hierarchy ends up being complicated by the need to shim in what is notably lacking in the geometric hierarchy. this leads to either an explosion of abstract classes, an explosion of object metadata (can this rectangle be stretched -- always 'no' for a square)? can it be scaled?), or a loss of the sort of functionality we typically want from inheritance. on the database level, these problems mostly don't exist in relational designs but they do exist in object relational designs, only differently (more on that below). the liskov substitution principle in application programming primarily operates primarily to preserve assumptions regarding state changes in an object when an interface is called. if a square (defined as a rectangle where x and y measurements are the same) can be stretched, it is no longer a square. on the other hand, if we can't stretch it, it may not fill the contract of the rectangle. when a square is-a rectangle subtype the square-rectangle problem can be turned on its head to some extent with the question of when this is-a relationship is valid. in fact the is-a relationship is valid for immutable objects. it is perfectly valid to have a class immutablesquare be a subclass of immutablerectangle. secondly the subtyping can be valid if sufficient information is attached to the class to specify the invariants. for example we might have a class with attributes "can_stretch" and "can_scale" where setting can_stretch off (as it would always be in a square) ensures that length to width proportions are always preserved. the problem here is that the subclass invariances require support in the superclass and this leads to a lot of complexity in implementing the superclass, as well as fundamental limits of what can then be subclassed. a look at database limitations in a normal relational database, the lsp is always met. domain-level constraints only apply to storage and not to calculation outputs, for example. in database systems where tables instantiate types and substitutability is not anticipated (oracle, db2), then the lsp applies in the same way it does in application programming. in the cases of table inheritance (informix, postgresql), things start to get a little weird. for the purpose of stored data, the lsp is generally satisfied because constraints are inherited and strengthened. a subtype is simply a sub-domain of the parent type possibly with some additional attributes. the square-rectangle problem in object-relational databases the square-rectangle problem is only even a potential problem in object-relational databases. purely relational designs can never run into this issue. in object-relational designs a few very specific issues can occur depending entirely on how object-relational functionality is implemented at the database level. in databases like oracle and db2, the lsp applies pretty much as is. in informix and postgresql, the problems faced are actually somewhat different and lead to new classes of anomilies that need to be considered. these include most prominently update anomilies when parent tables are updated. these do not necessarily have easy workarounds. consider the following table structure: create table my_rectangle ( id serial primary key, height numeric, width numeric ); create table my_square ( check (height = width) ) inherits (my_rectangle); now, we may insert a bunch of squares and rectangles, and we could have checks. moreover we could have custom triggers to verify referential integrity for rows referencing my_rectangle and my_square. so suppose we use update on my_rectangle to double the height of every rectangle in the system: update my_rectangle set height = height * 2; oops, we can't do that. while rows in my_rectangle will happily be adjusted, rows in the child table my_square will not, and you will get an error. this error can be avoided through a few issues, each of which has problems: disallow updates, only allow inserts and deletes. this makes referential integrity harder to manage when the row must be deleted and re-inserted. use a trigger to rewrite an update to be a delete plus insert into either the current or parent table depending on constraints. this makes ri harder to manage in that the procedure must disable custom ri triggers before doing this. this would require that the procedure be aware of all custom ri triggers in order to do this. conclusions the liskov substitution principle depends quite a bit on specifics of how an object-relational database system intersects tables and objects for how it applies. in general in a purely relational design you will never need to worry about it, but in an object-relational design there are some nasty corner cases that can come up, particularly where a query may operate on heterogenous subtypes as a set. this is, perhaps, the only one of the solid principles which is o-r specific and it hits the db in different ways than the application because the db operates on different principles.
February 11, 2013
by Chris Travers
· 4,808 Views
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Java 8: From PermGen to Metaspace
As you may be aware, the JDK 8 Early Access is now available for download. This allows Java developers to experiment with some of the new language and runtime features of Java 8. One of these features is the complete removal of the Permanent Generation (PermGen) space which has been announced by Oracle since the release of JDK 7. Interned strings, for example, have already been removed from the PermGen space since JDK 7. The JDK 8 release finalizes its decommissioning. This article will share the information that we found so far on the PermGen successor: Metaspace. We will also compare the runtime behavior of the HotSpot 1.7 vs. HotSpot 1.8 (b75) when executing a Java program “leaking” class metadata objects. The final specifications, tuning flags and documentation around Metaspace should be available once Java 8 is officially released. Metaspace: A new memory space is born The JDK 8 HotSpot JVM is now using native memory for the representation of class metadata and is called Metaspace; similar to the Oracle JRockit and IBM JVM's. The good news is that it means no more java.lang.OutOfMemoryError: PermGen space problems and no need for you to tune and monitor this memory space anymore…not so fast. While this change is invisible by default, we will show you next that you will still need to worry about the class metadata memory footprint. Please also keep in mind that this new feature does not magically eliminate class and classloader memory leaks. You will need to track down these problems using a different approach and by learning the new naming convention. I recommend that you read the PermGen removal summary and comments from Jon on this subject. In summary: PermGen space situation This memory space is completely removed. The PermSize and MaxPermSize JVM arguments are ignored and a warning is issued if present at start-up. Metaspace memory allocation model Most allocations for the class metadata are now allocated out of native memory. The klasses that were used to describe class metadata have been removed. Metaspace capacity By default class metadata allocation is limited by the amount of available native memory (capacity will of course depend if you use a 32-bit JVM vs. 64-bit along with OS virtual memory availability). A new flag is available (MaxMetaspaceSize), allowing you to limit the amount of native memory used for class metadata. If you don’t specify this flag, the Metaspace will dynamically re-size depending of the application demand at runtime. Metaspace garbage collection Garbage collection of the dead classes and classloaders is triggered once the class metadata usage reaches the “MaxMetaspaceSize”. Proper monitoring & tuning of the Metaspace will obviously be required in order to limit the frequency or delay of such garbage collections. Excessive Metaspace garbage collections may be a symptom of classes, classloaders memory leak or inadequate sizing for your application. Java heap space impact Some miscellaneous data has been moved to the Java heap space. This means you may observe an increase of the Java heap space following a future JDK 8 upgrade. Metaspace monitoring Metaspace usage is available from the HotSpot 1.8 verbose GC log output. Jstat & JVisualVM have not been updated at this point based on our testing with b75 and the old PermGen space references are still present. Enough theory now, let’s see this new memory space in action via our leaking Java program… PermGen vs. Metaspace runtime comparison In order to better understand the runtime behavior of the new Metaspace memory space, we created a class metadata leaking Java program. You can download the source here. The following scenarios will be tested: Run the Java program using JDK 1.7 in order to monitor & deplete the PermGen memory space set at 128 MB. Run the Java program using JDK 1.8 (b75) in order to monitor the dynamic increase and garbage collection of the new Metaspace memory space. Run the Java program using JDK 1.8 (b75) in order to simulate the depletion of the Metaspace by setting the MaxMetaspaceSize value at 128 MB. JDK 1.7 @64-bit – PermGen depletion Java program with 50K configured iterations Java heap space of 1024 MB Java PermGen space of 128 MB (-XX:MaxPermSize=128m) As you can see form JVisualVM, the PermGen depletion was reached after loading about 30K+ classes. We can also see this depletion from the program and GC output. Class metadata leak simulator Author: Pierre-Hugues Charbonneau http://javaeesupportpatterns.blogspot.com ERROR: java.lang.OutOfMemoryError: PermGen space Now let’s execute the program using the HotSpot JDK 1.8 JRE. JDK 1.8 @64-bit – Metaspace dynamic re-size Java program with 50K configured iterations Java heap space of 1024 MB Java Metaspace space: unbounded (default) As you can see from the verbose GC output, the JVM Metaspace did expand dynamically from 20 MB up to 328 MB of reserved native memory in order to honor the increased class metadata memory footprint from our Java program. We could also observe garbage collection events in the attempt by the JVM to destroy any dead class or classloader object. Since our Java program is leaking, the JVM had no choice but to dynamically expand the Metaspace memory space. The program was able to run its 50K of iterations with no OOM event and loaded 50K+ Classes. Let's move to our last testing scenario. JDK 1.8 @64-bit – Metaspace depletion Java program with 50K configured iterations Java heap space of 1024 MB Java Metaspace space: 128 MB (-XX:MaxMetaspaceSize=128m) As you can see form JVisualVM, the Metaspace depletion was reached after loading about 30K+ classes; very similar to the run with the JDK 1.7. We can also see this from the program and GC output. Another interesting observation is that the native memory footprint reserved was twice as much as the maximum size specified. This may indicate some opportunities to fine tune the Metaspace re-size policy, if possible, in order to avoid native memory waste. Now find below the Exception we got from the Java program output. Class metadata leak simulator Author: Pierre-Hugues Charbonneau http://javaeesupportpatterns.blogspot.com ERROR: java.lang.OutOfMemoryError: Metadata space Done! As expected, capping the Metaspace at 128 MB like we did for the baseline run with JDK 1.7 did not allow us to complete the 50K iterations of our program. A new OOM error was thrown by the JVM. The above OOM event was thrown by the JVM from the Metaspace following a memory allocation failure. #metaspace.cpp Final words I hope you appreciated this early analysis and experiment with the new Java 8 Metaspace. The current observations definitely indicate that proper monitoring & tuning will be required in order to stay away from problems such as excessive Metaspace GC or OOM conditions triggered from our last testing scenario. Future articles may include performance comparisons in order to identify potential performance improvements associated with this new feature. Please feel free to provide any comment.
February 11, 2013
by Pierre - Hugues Charbonneau
· 599,161 Views · 33 Likes
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How Do You Organise Maven Sub-Modules?
Being an itinerant programmer one of the things I've noticed over the years is that every project you come across seems to have a slightly different way of organising its Maven modules. There seems to be no conventional way of characterising the contents of a project's sub-modules and not that much discussion on it either. This is strange, as defining the responsibilities of your Maven modules seems to me to be as critical as good class design and coding technique to a project's success. So, in light of this dearth of wisdom, here's my two penneth worth... When you first come across a new project, you'll generally find a layout convention that vaguely matches that defined by the Better Builds With Maven manual. The 'clean' project directory usually contains a POM file, a src folder and several sub-modules, each in their own subdirectory, as shown in the diagram below: If we all agree that this is the standard way of approaching the top level of project layout (and I have seen it done slightly differently) then there seems to be three different approaches taken when organising the responsibilities of each of a project's sub-modules. These are: Totally haphazardly. By class type. By functional area. I'm not going to linger on those projects that are organised seemingly without any structure or order except to say that they probably started off well organised but were not designed well enough to endure the changes forced upon them. In saying that a project's sub-modules are organised 'by class type', I mean that modules are used to group together all classes that comprise, but are not limited to, a layer in the program's architecture. For example a module could contain all classes that make up the program's service or persistence layers or a module could contain all model class (i.e. beans). Conversely, in saying that a project's sub-modules are organised by functional area I'm talking about a situation where each module contains, as close as possible, a vertical slice of the application, including model beans, service layer, controllers etc. If the truth be told then there are any number of ways to organise your project's sub-modules. Most project set-ups are fairly flat in structure, which is what I've demonstrated above; however, if you take a look at Erik Putrycz's 2009 talk Maven – or how to automate java builds, tests and version management with open source tools, he demonstrates that you can have modules within modules within modules. In order to explore this a little further, I'm going to invent my usual preposterously contrived scenario and in this scenario, you've got to write a program for a Welsh dental practice owned by a man called Jones also known locally as 'Jones The Driller'. The requirements would be pretty standard, I suspect, for a dental practice and would include handling: Patients details: name, address, DOB, phone number etc. Medical records, including treatments and outcomes. Appointments. Accounting, e.g. sales, purchase, wages etc. Auditing: as in who did what to whom... As a solution to Jones The Driller's problem, you propose that you write a multi-module web application based upon Spring, MVC and tomcat that, when assembled, has a standard 'n' tier design of a mySQL database, a database layer, service layer, a set of controllers and some JSPs that comprise the view. In creating your project your idea is to organise your sub-modules 'by class type' and you come up with the following module organisation, shown below roughly in build dependency order dentists-model dentists-utils dentist-repository dentists-services dentists-controllers dentists-web ...which on your screen looks something like this: Your dentists-model module contains the project's beans that model object used from the persistence layer right up to the JSPs. dentists-repository, dentists-services and dentists-controllers reflect the various layers of your application, with dentists-web module containing all the JSPs, CSS and other view paraphernalia. As for dentists-utils, well every project has a utils module where all the really useful, but disparate classes end up. Meanwhile, in a different universe, a different version of you decides to organise your project's sub-modules by functional area and you come up with the following breakdown: dentists-utils dentists-audit dentists-user-details dentists-medical-records dentists-appointments dentists-accounts dentists-repository dentists-integration dentists-web In this scenario, the build order is somewhat different; virtually all modules will depend upon dentists-utils and, depending upon your exact audit requirements, most modules will rely upon dentists-audit. You can also see in the following images that the sub-module package structure has been arranged on layer and type boundaries in that each module has its own model, repository (which contains interface definitions only) services and controller packages and that the layout of each module is identical at the top level. Another discussion to have here is the organisation of your project's package structure, where you can ask the same kind of questions: do you organise 'by class type' or 'by functional area' as shown above? You may have noticed that the dentists-repository modules can be fairly near the end of the build cycle as it only contains the implementation of the repository classes and not their interface definitions. You may have also noticed that dentists-web is again a separate module. This is because you're a pretty savvy business guy and in keeping the JSPs etc. in their own module, you hope to re-skin your app and sell it to that other Welsh dentist down the road: Williams The Puller. From a test perspective, each module contains its own unit tests, but there's a separate integration test module that, as it'll take longer to execute can be run when required. There are generally two ways of defining integration tests: firstly by putting them in their own module, which I prefer, and secondly by using a integration test naming convention such as somefileIT.java, and running all *IT.java files separately to all *Test.java files. Your two identical selves have proposed two different solutions to the same problem, so I guess that it's now time to takes a look at the pros and cons of each. Taking the 'by class type' solution first, what can be said about it? On the plus side, it's pretty maintainable in that you always known where to find stuff. Need a service class? Then that's in the dentist-service module. Also, the build order is very straight forward. On the down side, organisation 'by class type' is prone to problems with circular dependencies creeping in and classes with totally different responsibilities are all mixed up together making it difficult to re-use functionality in other projects without unnecessarily dragging in the who shebang. So, what about the pros and cons of the 'by functional area' approach? To my way of thinking, given the package structure of each module, it's just about as easy to locate a class using this technique as it is when using 'by class type'. The real benefit of using this approach is that it's far simpler to re-use a functional area of code in other projects. For example, I've worked on many projects in different companies and have implemented auditing several times. Each time I implement it I usually do it in roughly the same way, so wouldn't it be good just to reuse the first implementation? Not withstanding code ownership issues... The same idea also applies to dentists-user-details; the requirement to manage names and addresses applies equally as well to a shoe sales web site as it does a dental practice. And the downside? One of the benefits of this approach is that the modules are highly decoupled, but from experience no matter how hard you try, you always end up with more coupling that you'd like. You may have already spotted that both of these proposals are not 100% pure; 'by class type' contains a bit of 'by functionality' and conversely 'by functional area' contains a couple of 'by class type' modules. This may be avoidable, but I'm purposely being pragmatic. As I said earlier you always see a utils module in a project. Furthermore creating a separate database module allows you to change your project's database implementation fairly easily, which may make testing easier in some circumstances and likewise, having a separate web module allows you to re-skin your code should you be lucky enough to sell the same product to multiple customers with their own branding. Finally, one of the unwritten truths in software development is that once you've organised your project into its sub-modules you'll rarely get the opportunity to reorganise and improve them: there usually isn't the time or the political will as doing so costs money; however, it should be remembered that, in Agile terms, project module composition is, like code, a form of technical debt, which if done badly also costs you a lot of cash. It therefore seems a really good idea, as a team, to plan out your project thoroughly before starting to code. So be radical, do some design or have a meeting, you know it'll be worth it in the end.
February 8, 2013
by Roger Hughes
· 28,809 Views · 1 Like
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Escaping Solr Query Characters In Python
I’ve been working in some Python Solr client code. One area where bugs have cropped up is in query terms that need to be escaped before passing to Solr. For example, how do you send Solr an argument term with a “:” in it? Or a “(“? It turns out that generally you just put a \ in front of the character you want to escape. So to search for “:” in the “funnychars” field, you would send q=funnychars:\:. Php programmer Mats Lindh has solved this pretty well, using str_replace. str_replace is a convenient, general-purpose string replacement function that lets you do batch string replacement. For example you can do: $matches = array("Mary","lamb","fleece"); $replacements = array("Harry","dog","fur"); str_replace($matches, $replacements,"Mary had a little lamb, its fleece was white as snow"); Python doesn’t quite have str_replace. There is translate which does single character to single character batch replacement. That can’t be used for escaping because the destination values are strings(ie “\:”), not single characters. There’s a general purpose “str_replace” drop-in replacement at this Stackoverflow question: edits =[("Mary","Harry"),("lamb","dog"),("fleece","fur")]# etc.for search, replace in edits: s = s.replace(search, replace) You’ll notice that this algorithm requires multiple passes through the string for search/replacement. This is because that earlier search/replaces may impact later search/replaces. For example, what if edits was this: edits =[("Mary","Harry"),("Harry","Larry"),("Larry","Barry")] First our str_replace will replace Mary with Harry in pass 1, then Harry with Larry in pass 2, etc. It turns out that escaping characters is a narrower string replacement case that can be done more efficiently without too much complication. The only character that one needs to worry about impacting other rules is escaping the \, as the other rules insert \ characters, we wouldn’t want them double escaped. Aside from this caveat, all the escaping rules can be processed from a single pass through the string which my solution below does, performing a heck of a lot faster: # These rules all independent, order of# escaping doesn't matter escapeRules ={'+':r'\+','-':r'\-','&':r'\&','|':r'\|','!':r'\!','(':r'\(',')':r'\)','{':r'\{','}':r'\}','[':r'\[',']':r'\]','^':r'\^','~':r'\~','*':r'\*','?':r'\?',':':r'\:','"':r'\"',';':r'\;',' ':r'\ '}defescapedSeq(term):""" Yield the next string based on the next character (either this char or escaped version """forcharin term:ifcharin escapeRules.keys():yield escapeRules[char]else:yieldchardefescapeSolrArg(term):""" Apply escaping to the passed in query terms escaping special characters like : , etc""" term = term.replace('\\',r'\\')# escape \ firstreturn"".join([nextStr for nextStr in escapedSeq(term)]) Aside from being a good general solution to this problem, in some basic benchmarks, this turns out to be about 5 orders of magnitude faster than doing it the naive way! Pretty cool, but you’ll probably rarely notice the difference. Nevertheless it could matter in specialized cases if you are automatically constructing and batching large/gnarly queries that require a lot of work to escape. Anyway, enjoy! I’d love to hear what you think!
February 8, 2013
by Doug Turnbull
· 4,927 Views
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Hbase Error: Region is not online: -ROOT-„0
If you are running HBase and commands are giving you an error that looks like this: Fri Oct 05 21:45:02 UTC 2012, org.apache.hadoop.hbase.client.ScannerCallable@74f2db2d, org.apache.hadoop.hbase.NotServingRegionException: org.apache.hadoop.hbase.NotServingRegionException: Region is not online: -ROOT-„0 at org.apache.hadoop.hbase.regionserver.HRegionServer.getRegion(HRegionServer.java:2859) at org.apache.hadoop.hbase.regionserver.HRegionServer.openScanner(HRegionServer.java:2071) at sun.reflect.NativeMethodAccessorImpl.invoke0(Native Method) at sun.reflect.NativeMethodAccessorImpl.invoke(NativeMethodAccessorImpl.java:39) at sun.reflect.DelegatingMethodAccessorImpl.invoke(DelegatingMethodAccessorImpl.java:25) at java.lang.reflect.Method.invoke(Method.java:597) at org.apache.hadoop.hbase.ipc.WritableRpcEngine$Server.call(WritableRpcEngine.java:364) at org.apache.hadoop.hbase.ipc. HBaseServer$Handler.run(HBaseServer.java:1345) Your HBase master is failing to start because Zookeeper is giving it an incorrect location for where the -ROOT- table is located. If you go into the HBase webui, you’ll see “Assigning ROOT region” as the status of the master node on startup. To fix this: $ hbase zkcli zookeeper_cli> rmr /hbase/root-region-server Then restart your master node, and you should be fine.
February 7, 2013
by George London
· 13,063 Views
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How to Compress and Uncompress a Java Byte Array Using Deflater/Inflater
Here is a small code snippet which shows an utility class that offers two methods to compress and extract a Java byte array.
February 6, 2013
by Ralf Quebbemann
· 135,356 Views · 2 Likes
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